AU593570B2 - Digital data processing system - Google Patents

Description

AUSTRALIA

Patents Act 593570 COMPLETE SPECIFICATION

(ORIGINAL)

Glass Int. Class Application Number: 6 15 'q Iu Lodged: n Lodged: 4 Accepted: O 1 jnin Complete Specificatio Published: Priorpty 94*~ 9 9 0*4'' 90 99 9 99 9 99 4 9 99* 099999 9 9 9 99 9 9 9944 99 99 4 9999 99 9 9 49 p t~ 4 9 *9 9.99 9 9.

94 1 9 j It S 9 C moey Related Art: APPLICANT'SREF,: USSN 266 401 et-al Div. of 60178/8 DATA GENERAL CORPORATION Name(s) of Applicant(s): Address(es) of Applicant(s): 4400 Computer Drive, Westboro, Massachusetts 01580, UNITED STATES OF AMERICA.

Actual Inventor(s): John K. Ahlstromr; Brett L. Bachman; David H. Bernstein; Richard Glenn Bratt; Gerald F. Clancy; Edward S. Gavrin; Ronald Hans Gruner; Lawrence H. Katz; Craig James Mundie; Michael S. Richmond; Stephen I. Schleimer; Steven J.

Wallach; Water A. Wallach, Jr.; Douglas, W0lls&; Richard A. Belgard; Thomas M. Jones PHILLIPS, ORMONDE AND FITZPATRICK Patent and Trade Mark Attorneys 367 Collins Street Melbourne, Australia, 3000 Address for Service is:, Complete Specification fo the invention entitled: "DIGITAL DATA PR~OCESSING SYSTEM" The following statement is a full description of this invention, including the best method of performing It known, to applicant(s):, P 19/3/84 DIGITAL DATA PROCESSING SYSTEM The present application is a voluntary division from Australian patent application no. 80178/82 lodged by the same applicant. Australian patent application no. 80178/82 is hereby incorporated by reference into the present application, The present invention relates to a digital data processing system and, more particularly, to a multiprocess digital data processing system in which system memory is organized into a plurality of objects, each object containing items of data, including instructions, and each object having associated with it a set of processing entities to which it will respond.

A general trend in the development of data processing systems has been towards systems suitable for use in inter-corinected data processing networks. Another trend has been towards data processing systems wherein the internal structure of the system is flexible, protected from users, and effectively invisible to the user and wherein the user is presented with a flexible and simplified interface to the system.

Certain problems and shortcomings affecting the realization of such a data processing system have appeared :repeatedly in the prior art and must be overcome to create a O data processing system having the above attributes. These prior art problems and limitations include the following topics.

.ip 1A 1 1 MJP 1A "i -2- First, the data processing systems of the prior art have not provided a system wide addressing system suitable for use in common by a large number of data processing systems interconnected into a network.

Addressing systems of the prior art have not provided sufficiently large address spaces and have not allowed information to be permanently and uniquely identified; Prior addressing systems have not made provisions for information to be located and identified as to type or 10 format, and have not provided sufficient granularity. In addition, prior addressing systems have reflected the physical structure of particular data processingsystems. That is, the addressing systems have been dependent upon whether a particular computer was, for example, an 8, 16, 32, 64 or 128 bit machine. Since prior data processing systems have incorporated addressing mechanisms wherein the actual physical structure of the processing system is apparent to the user, the operations a user could perform have been limited by the addressing mechanisms. In addition, prior processor systems have operated as fixed word length machines, further limiting user operations.

Prior data processing systems have not provided effective protection mechanisms preventing one user from I effecting another user's data and programs without permission. Such protection mechanisms have not, allowed unique, positive identification of users requesting i -3access to information, or of information, nor have such mechanisms been sufficiently flexible in operation. In addition, access rights have pertained to the users rather than to the information, so that control of access rights has been difficult. Finally, prior art protection mechanisms have allowed the use of "Trojan Horse arguments". That is, users not having access rights to certain information have been able to gain access to that information through another user or procedure having such access rights.

*Yet another problem of the prior art is that of providing a simple and flexible interface user interface ."Go to a data processing system. The character of user's Iinterface to a data processing system is determined, in part, by the means by which a user refers to and identifies operands and procedures of the user's programs l and by the instruction structure of the system. Operands and procedures are customarily referred to and identified by some form of logical address having points of 20 reference, and validity, only within a user's program.

These addresses must be translated into logical and physical addresses within a data processing system each S*time a program is executed, and must then be frequently S retranslated or generated during execution of a program.

25 In addition, a user must provide specific instructions as to data format and handling. As such reference to operands or procedures typically comprise a major portion i z r -4of the instruction stream of the user's program and requires numerous machine translations and operations to implement. A user's interface to a conventional system is thereby complicated, and the speed of execution of programs reduced, because of the complexity of. the program references to operands and procedures.

A data processing system's instruction structure includes both the instructions for controlling system operations and the means by which these instructions are 10 executed. Conventional data processing systems are 0 "designed to efficiently execute instructions in one or two user languages, for example, FORTRAN or COBOL.

Programs written in any other language are not efficiently executable. In addition, a user is often faced with difficult programming problems when using any high level language other than the particular one or two I languages that a particular conventional system is S" designed to utilize.

Yet another problem in conventional data processing systems is that of protecting the system's internal mechanisms, for example, stack mechanisms and internal control mechanisms, from accidental or malicious interference by a user.

Finally, the internal structure and operation of prior art data processing systems have not been flexible, or adaptive, in structure and operation. That is, the internal structure structure and operation of prior systems have not allowed the systems to be easily 1 modified or adapted to meet particular data processing requirements. Such modifications may include changes in internal memory capacity, such as the addition or deletion of special purpose subsystems, for example, floating point or array processors. In addition, such modifications have significantly effected the users interface with the system.

Ideally, the actual physical structure and operation of the data processing system should not be apparent at the user interface.

10 The present invention provides data processing system improvements and features which alleviate the above-described problems and limitations.

Specifically the present invention provides a digital computer system including: memory means for performing memory operations including storing and providing items of data, said items including instructions; t processor means connected to said memory means for

II

performing operations in response to said instructions 20 including processing said items of data; 4 4 e memory organization means operative on said memory means for organizing said memory means into objects identified by unique identifiers and allowing the location of said items therein, each said item being associated with an object and addressable by a logical address specifying I the unique identifier of the object with which said item is associated and an offset specifying the location of said item in its associated object, said memory organization Si means further identifying for each one of said objects a selected set of subjects representing entities for which said processor means responds to said instructions; memory operation specifier generation means in said processor means responsive to said instructions for providing a memory operation specifier for each item processed by said processor, said mv.mory operation specifier MJP -5 for a given item specifying the object in which said given item is to be located, a logical address for said given item, and one of said memory operations; and memory operations means responsive to a memory operation specifier for performing the memory operation specified by the memory operation specifier on the item specified by the memory operation specifier for a current subject for which said processor means is currently executing said instructions and which is one of the subjects in the Selected set of subjects for the object specified in said memory operation specifier.

The digital computer system of the present invention may include a memory system comprising mass storage devices and one or more processors connected to the memory system.

The memory system may be organized into objects containing data items, operands and instructions, each object *9*9 being identified by an object identifier. Locations of data 0 e items in the memory system may be specified by means of the object identifier for the object containing the data item.

The object identifier e o .4 S 9r t* e 99

I

4r MJP 6 I i in~*i~ am*nsp include a code field which identifies the time of creation of the object relative to a selected initial time which is common to a plurality of digital computer systems. The system further include a stack means which store the current state of execution of the system of the objects contain information relating to such current state of execution.

r t a a T a* 4.4 a a..

bat.

a.l 4. 43ija .4LL t"J~- Access to the data items in objects -e controlled by protection mechanisms of the systems. When the digital computer system processes data in an object, it d-e Sso for a subject representing an entity using the computer system.

m a ceI\e An access control list associated with each object definessets of subjects in a set of memory operations which a subject in a given set of subjects may perform on data items in the object. A memory operation on a data item in an object may sOcc<'eN ueetda only if there is an access control list entry associated with the object which allows the subject for whom the processor is performing the memory operation to perform that operation on the data in the object.

I| 4 i t t t I 4 4t '.ntf -7-

I-

-9level user languages.

It is a further object of e present invention to provide data processing i rnal mechanisms protected from user interferenc It is yet an ter object of the present invention to provide a d processing system having a flexible interna structure capable of multiple, concurrent LO Qther objct's, avantage nr fp atuirt nf t-htp.present *.coo invention will be understood by those of ordinary skill in the art, after referring to the following detailed S description of the preferred embodiments and drawings wherein: 6 Fig. 1 is a partial block diagram of a computer S system incorporating the present invention; Fig. 2 is a diagram illustrating computer system addressing structure of the present invention; Fig. 3 is a diagram illustrating the computer system instruction stream of the present invention; Fig. 4 is a diagram illustrating the control S structure of a conventional computer system; Fig. 4A is a diagram illustrating the control structure of a computer system incorporating the present invention; Fig. 5 Fig. Al inclusive are diagrams all relating to the present invention;

VAI

Fig. 5 is a diagram illustrating a stack mechanism; Fig. 6 is a diagram illustrating procedures, procedure objects, processes, and virtual processors; Fig. 7 is a diagram illustrating operating levels and mechanisms of the present computer; Fig. 8 is a diagram illustrating a physical implementation of processes and virtual processors; Fig. 9 is a diagram illustrating a process and 10 process stack objects; Fig. 10 is a diagram illustrating operation of macrostacks and secure stacks; Fig. 11 is a diagram illustrating detailed structure of a stack; Fig. 12 is a diagram illustrating a physical descriptor; ir Fig. 13 is a diagram illustrating the relationship between logical pages and frames in a memory storage space; 20 Fig. 14 is a diagram illustrating access control to objects; Fig. 15 is a diagram illustrating virtual processors and virtual processor swapping; Fig. 16 is a partial block diagram of an I/O system of the present computer system; Fig. 17 is a diagram illustrating operation of a ring grant generator;

I.

-11- Fig. 18 is a partial block diagram of a memory system; Fig. 19 is a partial block diagram of a fetch unit of the present computer system; Fig. 20 is a partial block diagram of an execute unit of the present computer system; Fig. 101 is a more detailed partial block diagram of the present computer system; Fig. 102 is a diagram illustrating certain t..

10 information structures and mechanisms of the present computer system; Fig. 103 is a diagram illustrating process structures; Fig. 104 is a diagram illustrating a macrostack structure; Fig. 105 is a diagram illustrating a secure stack structure; Figs. 106 A, B, and C are diagrams illustrating the addressing structure of the present computer system; Fig. 107 is a diagram illustratig addressing mechanisms of the present computer system; Fig. 108 is a diagram illultrating a name table S* entry; Fig. 109 is a diagram illustrating protection mechanisms of the present computer system; Fig. 110 is a diagram illustrating instruction and microinstruction mechanism of the present computer system; -12- Fig. 201 is a detailed block diagram of a memory system; Fig. 202 is a detailed block diagram of a fetch unit; Fig. 203 is a detailed block diagram of an execute unit; Fig, 204 is a detailed block diagram of an I/O system; Fig. 205 is a partial block diagram of a 10 diagnostic processor system; Fig. 206'is a diagram illustrating assembly of Figs. 201-205 to form a detailed block diagram of the S present computer system; S. 1 Fig. 207 is a detailed block diagram of a memory 15 interface controller; Fig. 209 is a diagram of a memory to I/O system S port interface; Fig. 210 is a diagram of a memory operand port interface; S. 20 Fig. 211 is a diagram of a memory instruction port interface; Fig. 230 is a detailed block diagram of memory field interface unit logic; S :i Fig. 231 is a diagram illustrating memory format manipulation operations; Fig. 238 is a detailed block diagram of fetch unit offset multiplexer; 1 t 1 1

I

-13- Fig. 239 is a detailed block diagram of fetch unit bias logic; Fig. 240 is a detailed block diagram of a generalized four way, set associative cache representing name cache, protection cache, and address translation unit; Fig. 241 is a detailed block diagram of portions of computer system instruction and microinstruction control logic; 10 Fig. 242 is a detailed block diagram of portions of computer system microinstruction control logic; .o Fig. 243 is a detailed block diagram of further portions of computer system microinstruction control C logic; Fig. 244 is a diagram illustrating computer system states of operation; t, Fig. 245 is a diagram illustrating computer t system states of operation for a trace trap request; Fig. 246 is a diagram illustrating computer f' 20 system states of operation for a memory repeat interrupt; Fig. 247 is a diagram illustrating priority level and masking of computer system events; Fig. 248 is a detailed block diagram of event logic; Fig. 249 is a detailed block diagram of microinstruction control store logic; Fig. 251 is a diagram illustrating a return control word stack word; ii

I

I

-14- 9.99 9 9 9 9 *9 9 9*99.9 .9 9 9999 99 9 9 99 ii e 9 .o 9 Fig. 252 is a diagram illustrating machine control words; Fig. 253 is a detailed block diagram of a register address generator; Fig. 254 is a block diagram of interval and egg timers; Fig. 255 is a detailed block diagram of execute unit control logic; Fig. 257 is a detailed block diagram of execute 10 unit multiplier data paths and memory; Fig. 260 is a diagram illustrating operation of an execute unit command queue load and interface to a fetch unit; Fig. 261 is a diagram illustrating operation of an execute unit operand buffer load and interface to a fetch unit; Fig. 262 is a diagram illustrating operation of an execute unit storeback or transfer of results and interface to a fetch unit; 20 Fig. 263 is a diagram illustrating operation of an execute unit check test condition and interface to a fetch unit; Fig. 264 is a diagram illustrating operation of an execute unit exception test and interface to a fetch unit; Fig. 265 is a block diagram of an execute unit arithmetic operation stack mechanism; *999 99 9 1 *49999 Ai 1 .^r i r i.iL i 1(1 4i 44c 4 *c 4 4 4c 48 *44 .4 4 4 .4.

*c 4 4 44: Fig. 266 is a diagram illustrating execute unit and fetch unit interrupt handshaking and interface; Fig. 267 is a diagram illustrating execute unit and fetch unit interface and operation for nested interrupts; Fig. 268 is a diagram illustrating execute unit and fetch unit interface and operation for loading an execute unit control store; Fig. 269 is a detailed block diagram and 10 illustration of operation of an I/O system ring grant generator; Fig. 270 is a detailed block diagram of a fetch unit micromachine of the present computer system; Fig. 271 is a diagram illustrating a logical 15 descriptor; Fig. 272 is a diagram illustrating use of fetch unit stack registers; Fig. 273 is a diagram ilustrating structures controlling event invocations; Fig. 301 is a diagram illustrating pointer formats; Fig. 302 is a diagram illustrating an associated address table; Fig. 303 is a diagram illustrating a namespace overview of a procedure object; Fig. 304 is a diagram illustrating name table entries; 44.4 4* 44 4 4 S 4 "1 i 1 -16- Fig. 305 is a diagram illustrating an example of name resolution; Fig. 306 is a diagram illustrating name cache entries; Fig. 307 is a diagram illustrating translation of S-interpreter universal identifiers to dialect numbers; Fig. 401 is a diagram illustrating operating systems and system resources; Fig. 402 is a diagram illustrating multiprocess operating systems; ~Fig. 403 is a diagram illustrating an extended operating system and a kernel operating system; Fig. 404 is a diagram illustrating an EOS view of objects; Fig. 405 is a diagram illustrating pathnames to universal identifier translation; Fig. 406 is a diagram illustrating universal identifier detail; 20 Fig. 407 is a diagram illustrating address translation with an address translation unit, a memory hash table, and a memory; Fig. 408 is a diagram illustrating hashing in an s 'active subject table; Fig. 409 is a diagram illustrating logical allocation units and objects; i i

A

r I 7 -17- Fig. 410 is a diagram illustrating an active allocation unit table and active allocation logical units; a U 4 *i *e 4 ai a a .4 a a a1 a a a *4.a r* 44 *4 a 4 49r a.

a. a 9 44 a. a; a Fig. 411 is a diagram illustrating a conceptual logical allocation unit directory structure; Fig. 412 is a diagram illustrating detail of a logical allocation unit directory entry; Fig. 413 is a diagram illustrating universal identifiers and active object numbers; 10 Fig. 416 is a diagram illustrating subject templates, primitive access control list entries, and extended access control list entries; Fig. 421 is a diagram illustrating an active primitive access matrix and an'active primitive access matrix entry; i Fig. 422 is a diagram data access checking; Fig. 448 is a diagram counters and await entries; 20 Fig. 449 is a diagram table overview; Fig. 453 is a diagram of a virtual processor; Fig. 454 is a diagram processor synchronization; Fig. 467 is a diagram of a macrostack object; illustrating primitive illustrating event illustrating an await illustrating an overview illustrating virtual illustrating an overview 1~ ri

LI

f A -18- Fig. 468 is a diagram macrostack object base; Fig. 469 is a diagram maceostack frame; Fig. 470 is a diagram of a secure stack; Fig. 471 is a diagram secure stack frame; and, Fig. 472 is a diagram 10 of procedure object.

illustrating details of a illustrating details of a illustrating an overview illustrating details of a illustrating an overview 9.9, 49 c9 99 9c *l .9 9i 9 9* U 4 9 9.

.4 .9 9 i 94 9 '9 9 99

SD

The following description presents the structure and operation of a computer system incorporating a presently preferred embodiment of the present invention. As 15 indicated in the following Table of Contents, certain features of computer system structure and operation will first be described in an Introductory Overview. Next, these and other features will be described in further detail in a more detailed Introduction to the detailed 20 descriptions of the computer system. Following the Introduction, the structure and operation of the computer system will be described in detail. The detailed descriptions will present descriptions of the structure and operation of each of the major subsystems, or elements, of the computer system, of the interfaces between these major subsystems, and of overall computer system operation. Next, certain features of the operation of the individual subsystems will be presented in further detail.

.1 i!" Br

V

4- -19- Certain conventions are used throughout the following .escriptions to enhance clarity of presentation. First, and with exception of the Introductory Overview, each figure referred to in the following descriptions will be referred to by a three digit number. The most significant digit represents the number of the chapter in the following descriptions in which a particular figure is first referred to. The two least significant digits represent the sequential number of appearance of a figure 10 in a particular chapter. For example, Figure 319 would be the nineteenth figure appearing in the third chapter.

Figures appearing in the Introductory Overview are S referred to by a one or two digit number representing the order in which they are referred to in the Introductory 15 Overview. It should be noted that certain figure numbers, for example, Figure 208, do not appear in the Sfollowing figures and descriptions; the subject matter of these figures has been incorporated into other figures and these figures deleted, during drafting of the following descriptions, to enhance clarity of presentation.

Second, reference numerals comprise a two digit S* number (00-99) preceded by the number of the figure in m which the corresponding elements first appear. For example, reference numerals 31901 to 31999 would refer to elements 1 through 99 appearing in Fig. 319.

ii I 4.| Finally, interconnections between related circuitry is represented in two ways. First, to enhance clarity of presentation, interconnections between circuitry may be represented by common signal names or references, rather than by drawn representations of wires or buses. Second, where related circuitry is shown in two or more figures, the figures may share a common figure number and will be distinguished by a letter designation, for example, Figs.

319, 319A, and 319B. Common electrical points between o such circuitry may be indicated by a bracket enclosing a lead to such a point and a designation of the form indicates other figures having the same common point I for example, 319A, and designates the particular I common electrical point. In cases of related circuitry shown in this manner in two or more figures, reference numerals to elements will be assigned in sequence through I the group of figures; the figure number portion of such S, reference numerals will be that of the first figure of the group of figures.

INTRODUCTORY-

OVERVIEW

A. Hardware Overview (Fig. 1) B. Individual Operating Features (Figs. 2, 3 4, 5, 6) t* 1. Addressing (Fig. 2) 2. S-Language Instructions and Namespace Addressing (Fig. 3) 3. Architectural Base Pointer Addressing 4. Stack Mechanisms (Figs. r rf -21- .c

I

*4 t *i I Si S S S 45*.

*c I

S.

C. Procedure Processes and Virtual Processors (Fig. 6) D. CS 101 Overall Structure and Operation (Figs. 7, 8, 9, 10, 11, 12, 13, 14, 1. Introduction (Fig. 7) 2. Compilers 702 (Fig. 7) 3. Binder 703 (Fig. 7) 4. EOS 704 (Fig. 7) KOS and Architectural Interface 708 (Fig. 7) 6. Processes 610 and Virtual Processors 612 (Fig.

8) 7. Processes 610 and Stacks (Fig. 9) 8. Processes 610 and Calls (Figs. 10, 11) 9. Memory References and the Virtual Memory Management System (Fig. 12, 13) 10. Access Control (Fig. 14) 11. Virtual Processors and Virtual Processor Swapping (Fig. E. CS 101 Structural Implementation (Figs. 16, 17, 18, 19, 1. (IOS) 116 (Figs. 16, 17) 2. Memory (MEM) 112 (Fig. 18) 3. Fetch Unit (FU) 120 (Fig. 19) 4. Execute Unit (EU) 122 (Fig. 1. 1 ntiodu -itl on- '(igs 4-0-1+0j A. General Structure and Operation (Fig. 101) a. General Structure b. General Operation i:e I i II i

~I

i: :j c i L~ i i -22-

U

4* 4*

U,

UU

U U 5~ 44 0r c. Definition of Certain Terms d. Multi-Program Operation e. Multi-Language Operation f. Addressing Structure g. Protection Mechanism B. Computer System 10110 Information Structure and Mechanisms (Figs. 102, 103, 104, 105) a. Introduction (Fig. 102) b. Process Structures 10210 (Figs. 103, 104, 10 105) 1. Procedure Objects (Fig. 103) 2. Stack Mechanisms (Figs. 104, 105) 3. FURSM 10214 (Fig. 103) C. Virtual Processor State Blocks and Virtual 15 Process Creation (Fig. 102) D. Addressing Structures'10220 (Figs. 103, 106, 107, 108) 1. Objects, UID's, AON's, Names, and Physical Addresses (Fig. 106) 20 2. Addressing Mechanisms 10220 (Fig. 107) 3. Name Resolution (Figs. 103, 108) 4. Evaluation of AON Addresses to Physical Addresses (Fig. 107) E. CS 10110 Protection Mechanisms (Fig. 109) F. CS 10110 Micro-Instruction Mechanisms (Fig. 110) G. Summary of Certain CS 10110 Features and Alternate Embodiments.

~I

-23- 2. Dtale-D ai- ese tion-o-e--S 1110 Ma or- ubstms Fis. .201-206-.--20-7--274i A. MEM 10110 (Figs. 201, 206, 207-237) a. Terminology b. MEM 10112 Physical Structure (Fig. 201) c. MEM 10112 General Operation d. MEM 10112 Port Structure 1. 10 Port Characteristics 2. JO Port Characteristics 10 3. JI Port Characteristics e. MEM 10112 Control Structure and Operation (Fig. 207) SI 1. MEM 10112 Control Structure 2. MEM 10112 Control Operation f. MEM 10112 Operations g. MEM 10112 interfaces to JP 10114 and IOS tit 10116 (Figs. 209, 210, 211, 204) 1. 10 Port' 20910 Operating S. Characteristics (Figs 209, 204) 2. JO Port 21010 Operating Characteristics (Fig. 210) I JI Port 21110 Operating Charactertiscis (Fig. 211) h. FIU 20120 (Figs. 201, 230, 231) B. Fetch Unit 10120 (Figs. 202, 206, 101, 103, 104, 238) 1. Descriptor Processor 20210 (Figs. 202, 101, 103, 104, 238, 239) l i 1 f 1 r -24ereQ

C

a. Offset Processor 20218 Structure b. AON Processor 20216 Structure c. Length Processor 20220 Structure d. Descriptor Processor 20218 Operation a.a. Offset Selector 20238 b.b. Offset Multiplexer 20240 Detailed Structure (Fig. 238) c.c. Offset Multiplexer 20240 Detailed Operation aaa. Internal Operation bbb. Operation Relative to DESP 20210 e. Length Processor 20220 (Fig. 239) a.a. Length ALU 20252 b.b. BIAS 20246 (Fig. 239) f. AON Processor 20216 a.a. AONGRF 20232 b.b. AON Selector 20248 2. Memory Interface 20212 (Figs. 106, 240) a.a. Descriptor Trap 20256 and Data Trap 20258 b.b Name Cache 10226, Address Translation Unit 10228, and Protection Cache 10234 (Fig. 106) c.c. Structure and Operation of Generalized Cache and NC 10226 (Fig. 240) d.d. ATU 10228 and PC 10234 3. Fetch Unit Control Logic 20214 (Fig. 202) a.a. Fetch Unit Control Logic 20214 Overall Structure

C.

C

C C C CC C C C CC

CC..

Cr C I j r

'A

~-zF~ b.b. Fetch Unit Control Logic 20214 Operation a.a.a. Prefetcher 20264, Instruction Buffer 20262, Parser 20264, Operation Code Register 20268, CPC 20270, IPC 20272, and EPC 20274 (Fig. 241) Fetch Unit Dispatch Table 11010, Execute Unit Dispatch Table 20266, and Operation Code Register 20268 (Fig.

9 9 *4**l .4 9. 4 #4f 44 t #4 44 4 b'b.b.

242) c.c.c. Next Address Generator 24310 (Fig. 243) cc. FUCTL 20214 Circuitry for CS 10110 Internal Mechanisms (Figs. 244-250) a.a.a. State Logic 20294 (Figs.

244A-244Z) b.b.b. Event Logic 20284 (Figs.

245, 246, 247, 248) c.c.c. Fetch Unit S-Interpreter Table 11012 (Fig. 249) d.d. CS 10110 Internal Mechanism Control a.a.a. Return Control Word Stack 10358 (Fig. 251) b.b.b. Machine Control Block (Fig.

252) I 8r 'n I

I

-26c.c.c. Register Address Generator 20288 (Fig. 253) d.d.d. Timers 20296 (Fig. 254) e.e.e. Fetch Unit 10120 Interface to Execute Unit 10122 C. Execute Unit 10122 (Figs. 203, 255-268) a. General Structure of Execute Unit 10122 1. Execute Unit I/O 20312 Execute Unit Control Logic 20310 10 3. Multiplier Logic 20314 SV" 4. Exponent Logic 20316 5. Multiplier Control 20318 6. Test and Interface Logic 20320 b. Execute Unit 10122 Operation (Fig. 255) 1. Execute Unit Control Logic 20310 (Fig, 255) a.a. Command Queue 20342 b.b. Command Queue Event Control Store 25514 and Command Queue Event Address Control Store 25516 c.c. Execute Unit S-Interpreter Table 20344 d.d. Microcode Control Decode Register 20346 e.e. Next Addrefs Generator 20340 2. Operand Buffer 20322 (Fig. 256) 3. Multiplier 20314 (Figs. 257, 258) a 2. Multiplier 20314 I/O Data Paths and Memory (Fig 257) 4 y 4 1 r -27- 9*99 9 4r 4 .9 *r 9 9r 99 9 9:~ a 9 41i 4 1 1 41 *L 9 99 9 95 .9.4 49 t9 9 9i 1. 4 a.a.a. Container Size Check b.b.b. Final Result Output Multiplexer 20324 4. Test and Interface Logic 20320 (Figs.

260-268) a.a. FU 10120/EU 10122 Interface Loading of Command Queue 20342 (Fig. 260) b.b.b. Loading of Operand Buffer 20320 (Fig. 261) c.c.c. Storeback (Fig. 262) d.d.d. Test Conditions (Fig.

263) e.e.e. Exception Checking (Fig. 264) f.f.f. Idle Routine g.g.g. EU 10122 Stack Mechanisms (Figs. 265, 266, 267) h.h.h. Loading of Execute Unit S-Interpreter Table 20344 (Fig. 268) System 10116 (Figs. '204, 206, 269) I/O System 10116 Structure (Fig. 204) I/O System 10116 Operation (Fig. 269) 1. Data Channel Devices 2. I/O Control Processor 20412 3. Data Mover 20410 (Fig. 269) a.a. Input Data Buffer 20440 and Output Data Buffer 20442 D. I/O a.

b.

r* "i I -28b.b. Priority Resolution and Control 20444 (Fig. 269) E. Diagnostic Processor 10118 (Fig. 101, 205) F. CS 10110 Micromachine Structure and Operation (Figs. 270-274) a. Introduction b. Overview of Devices Comprising FU Micromachine (Fig. 270) 1. Devices Used By Most Microcode R 4 A a.a. MOD Bus 10144, JPD Bus 10142, and SDB Bus 27021 b.b. Microcode Addressing Sc.c.. Descriptor Processor 20218 (Fig.

271) d.d. EU 10122 Interface 2. Specialized Micromachine Devices I a.a. Instruction Stream Reader 27001 1 i 4 b.b. SOP Decoder 27003 c.c. Name Translation Unit 27015 d.d. Memory Reference Unit 27017 e.e. Protection Unit 27019 f.f. KOS Micromachine Devices c. Micromachine Stacks and Microroutine Calls SI and Returns (Figs. 272, 273) 1. Micromachine Stacks (Fig. 272) 2. Micromachine Invocations and Returns 3. Means of Invoking Microroutines i i 4. Occurrence of Event Invocations (Fig.

273) d. Virtual Micromachines and Monitor Micromachine 1. Virtual Mode 2. Monitor Micromachine e. Interrupt and Fault Handling 1. General Principles 2. Hardware Interrupt and Fault Handling 10 in CS 10110 3. Monitor Mode: Differential Masking a-29and Hardware Interrupt Handling g. FU Micromachine and CS 10110 Subsystems 3. Micromachine irtual Mode...

eA. Pointers and Pointer Resolution (Figs. 301 302) a. Pointer Formats (Fig. 301) b. Pointers in FU 10120 (Fig. 302) c. Descriptor to Pointer Conversion B. Namespace and the S-Interpreters (Figs. 303-307) a. Procedure Object 606 Overview (Fig. 303) b. Namespace 1. Name Resolution and Evaluation 2. The Name Table (Fig. 304) 3. Architectural Base Pointers (Figs.

305, 306) F rI;: a.a. Resolving and Evaluating Names (Fig. 305) b.b. Implementation of Name Evaluation and Name Resolve in CS 10110 c.c. Name Cache 10226 Entries (Fig.

306) d.d. Name Cache 10226 Hits e.e. Name Cache 10226 Misses f.f. Flushing Name Cache 10226 g.g. Fetching the Instruction Stream h.h. Parsing the Instruction Stream c. The S-Interpreters (Fig. 307) 1. Translating SIP into a Dialect Number (Fig. 307) 2. Dispatching 0* S riac

I

15 4. The- Kerne-l- peratin--S-stem A. Introduction S, a. Operating Systems (Fig. 401) 1. Resources Controlled by Operating Systems (Fig. 402) 20 b. The Operating System in CS 10110 c. Extended Operating System and the Kernel Operating System (Fig. 403) B. Objects and Object Management (Fig. 404) a. Objects and User Programs (Fig. 405) b. UIDs 40401 (Fig. 406) c. Object Attributes 'i if ni 3 -31- 9 *9#9 #9 a, a 4* a.

a I I ta~a~ a 4 1£ a I

I

a Ia a, I as., 4* 9 9* 44 9 4* 4.

tell I t I 4,

I

111111 C C Attributes and Access Control Implementation of objects 1. Introduction (Figs 407, 408) 2. Objects in Secondary Storage 10124 (Figs 409, 410) a.a. Representation of an Object's Contents on Secondary Storage 10124 b.b. LAUD 40903 (Figs. 411, 412) 3. Active Objects (Fig. 413) a.a. UID 40401 to AON 41304 Translation Access Control System- Subj ects Domains Access Control Lists 1. Subject Templates (Fig. 416) 2. Primitive Access Control Lists (PACLs) 3. APAM 10918 and Protection Cache 10234 (Fig. 421) 4. Protection Cache 10234 and Protection Checking (Fig. 422) C. The a.

b.

c.

4-, ,1 Processes 1. Synchronization of Processes 610 and Virtual Processors 612 a. Event Counters 44801, Await Entries 44804, and Await Tables (Fig. 448, 449) -32b. Synchronization with Event Counters 44801 and Await Entries 44804 E. Virtual Processors 612 (Fig. 453) a. Virtual Processor Management (Fig. 453) b. Virtual Processors 612 and Synchronization (Fig. 454) F. Process 610 Stack Manipulation 1. Introduction to Call and Return 2. Macrostacks (MAS) 502 (Fig. 467) a.a. MAS Base 10410 (Fig. 468) b.b. Per-domain Data Area 46853 (Fig.

468) c.c. MAS Frame 46709 Detail (Fig. 469) 3. SS 504 (Fig. 470) S 15 a.a. SS Base 47001 (Fig. 471) SS Frames 47003 (Fig 471) a.a.a. Ordinary SS Frame Headers 1051.4 (Fig.

471) b.b.b. Detailed Structure of Macrostate 10516 (Fig.

471 4.

c.c.c. Cross-domain SS Frames 47039 (Fig. 471) t 25 4. Portions of Procedure Object 608 t Relevant to Call and Return (Fig. 472) Execution of Mediated Calls a.a. Mediated Call SINs

LJ

h ;i 1_ -"n

I

-33b.b. Simple Mediated Calls (Figs. 270, 468, 469, 470, 471, 472) c.c. Invocations of Procedures 602 Requiring SEBs 46864 (Figs. 270, 468, 469, 470, 471, 472) d.d. Cross-Procedure Object Calls (Figs. 270, 468, 469, 470, 471, 472) e.e. Cross-Domain Calls (Figs. 270, 408, 418, 468, 469, 470, 471, 472) f.f. Failed Cross-Domain Calls (Figs.

270, 468, 469, 470, 471, 472) 6. Neighborhood Calls (Figs. 468, 469, 472) 'Set t 'it i tt *t I

LI

i 44 4 5 4r 4* 4 50L1CF 4 4 4.

4iiti

C

TNTRODUCTORY- OVERVIEW The following overview will first briefly describe the overalll physical structure and operation of a presently preferred embodiment of a digital computer 20 system incorporating the present invention. Then certain operating features of that computer system will be individually described. Next, overall operation of the computer system will be described in terms of those individual features.

S3 I ii: ii i Mr-, -34- A. Hardware- Overview-- Referring to Fig. 1, a block diagram of Computer System (CS) 101 incorporating the present invention is shown. Major elements of CS 101 are I/O System (IOS) 116, Memory (MEM) 112, and Job Processor (JP) 114. JP 114 is comprised of a Fetch Unit (FU) 120 and an Execute Unit (EU) 122. CS 101 may also include a Diagnostic Processor not shown or described in the instant description.

Referring first to IOS 116, a primary'function of IOS 116 is control of transfer of information between MEM 112 and the outside world. Information is transferred from MEM 112 to IOS 116 through IOM Bus 130, and from IOS 116 4 to MEM 112 through MIO Bus 129. IOMC Bus 131 is 4i 6 S 15 comprised of bi-directional control signals coordinating 'operation of MEM 112 and IOS 116. IOS 116 also has an S interface to FU 120 through IOJP Bus 132. IOJP Bus 132 is a bi-directional control bus comprised essentially of two interrupt lines. These interrupt lines allow FU 120 to indicate to IOS 116 that a request for information by FU 120 has been placed in MEM 112, and allows IOS 116 to inform FU 120 that information requested by FU 120 has been transferred into a location in MEM 112. MEM 112 is CS 101's main memory and serves as the path for 25 information transfer between the outside world and JP S114. MEM 112 provides instructions and data to FU 120 and EU 122 through Memory Output Data (MOD) Bus 140 and receives information from FU 120 and EU 122 through Job JP 114 is CS 101's CPU and, as described above, is comprised of FU 120 and EU 122. A primary function of FU 120 is executing operations of user's programs. As part of this function, FU 120 controls transfer of instructions and data from MEM 112 and transfer of results of JP 114 operations back to MEM 112. FU 120 also performs operating system type functions, and is capable of operating as a complete, general purpose CPU.

r EU 122 is primarily an arithmetic and logic unit provided to relieve FU 120 of certain arithmetic operations. FU 120, however, is capable of performing EU 122 15 operations. In alternate embodiments of CS 101, EU 122 may be provided only as an option for users having particular arithmetic requirements. Coordination of FU 120 and EU 122 operations is accomplished through FU/EU (FUEU) Bus 148, which includes bi-directional control signals and mutual interrupt lines. As described further' below, both FU 120 and EU 122 contain register file Sto arrays referred to respectively as CRF and ERF, in addition to registers associated with, for example, ALUs.

A primary feature of CS 101 is that IOS 116, MEM 112, 25 FU 120 and EU 122 each contain separate and independent microinstruction control, so that IOS 116, MEM 112, and S EU 122 operate asynchronously under the general control J of FU 120. EU 122, for example, may execute a complex :r 1 -36arithmetic operation upon receipt of data and a single, initial command from FU 120.

Having briefly described the overall structure and operation of CS 101, certain features of CS 101 will be individually further described next below.

B. Individual -Oeratin--Peatues-- Pasi-2- 6 1. Addressing (Fi 2-) Referring to Fig. 2, a diagramic representation of portions of CS 101's addressing structure is shown. CS 101's addressing structure is based upon the concept of Objects. An Object may be regarded as a container for holding a particular type of information. For example, one type of Object may contain data while another type of Object may contain instructions or procedures, such as a user program. Still another type of Object may contain microcode. In general, a particular Object may contain only one type or class of informtion. An Object may, for example, contain up to 2 bits of information, but the actual size of a particular Object is flexible. That is, the actual size of a particular Object will increase as information is written into that Object and will decrease as information is taken from that Object. In general, information in Objects is stored sequentially, that is without gaps.

25 Each Object which can ever exist in any CS 101 system S is uniquely identified by a serial number referred to as a Unique Identifier (UID). A UID is a 128 bit value comprised of a serial number dependent upon, for example, i -37the particular CS 101 system and user, and a time code indicating time of creation of that Object. UIDs are permanently assigned to Objects, no two Objects may have the same UID, and UIDs may not be reused. UIDs provide an addressing base common to all CS 101 systems which may ever exist, through which any Object ever created may be permanently and uniquely identified.

As described above, UIDs are 128 bit values and are thus larger than may be conveniently handled in present embodiments of CS 101. In each CS 101, therefore, those Objects which are active (currently being used) in that system are assigned 14 bit Active Object Numbers (AONs).

Each Object active in that system will have a unique AON. Unlike UIDs, AONs are only temporarily assigned to 15 particular Objects. AONs are valid only within a particular CS 101 and are not unique between systems. An Object need not physically reside in a system to be assigned an AON, but can be active in that system only if it has been assigned an AON.

A particular bit within a particular Object may be identified by means of a UID address or an AON address.

In CS 101, AONs and AON addresses are valid only within t JP 114 while UIDs and UID addresses are used in MEM 112 and elsewhere. UID and AON addresses are formed by 25 appending a 32 bit Offset field to that Object's UID or AON. 0 fields indicate offset, or location, of a particular bit relative to the start of a particular Object.

F I ;Q X -38- Segments of information (sequences of information bits) within particular Objects may be identified by means of descriptors. A UID descriptor is formed by appending a 32 bit Length field of a UID address. An AON, or logical descriptor is formed by appending a 32 bit L field to an AON address. L fields identify length of a segment of information bits within an Object, starting from the information bit identified by the UID or AON address. In addition to length information, UID and logical descriptors also contain Type fields containing information regarding certain characteristics of the information in the information segment. Again, AON based descriptors are used within JP 114, while UID f *based descriptors are used in MEM 112.

Referring to Figs. 1 and 2 together, translation between UID addresses and descriptors and AON addresses and descriptors is performed at the interface between MEM 112 and JP 114. That is, addresses and descriptors within JP 114 are in AON form while addresses and descriptors in MEM 112, IOS 116, and the external world are in UID form. In other embodiments of CS 101 using AONs, transformation from UID to AON addressing may occur at other interfaces, for example at the IOS 116 to MEM S112 interface, or at the IOS 116 to external world 25 interface. Other embodiments of CS 101 may use UIDs i throughout, thae is not use AONs even in JP 114.

i -39- Finally, information within MEM 112 is located through MEM 112 Physical Addresses identifying particular physical locations within MEM 112's memory space. Both IOS 116 and JP 114 address information within MEM 112 by providing physical addresses to MEM 112. In the case of physical addresses provided by JP 114, these addresses are referred to as Physical Descriptors (PDs). As described below, JP 114 contains circuitry to translate logical descriptors into physical descriptors.

2. S-Language-' nstructions-and--Namespace Addressing Fia.--31 Pmor CS 101 is both an S-Language machine and a Namespace S machine. That is, operations to be executed by CS 101 are expressed as S-Language Operations (SOPs) while ft a 15 operands are identified by Names. SOPs are of a lower, So more detailed, level than user language instructions, for .S example FORTRAN and COBOL, but of a higher level than conventional machine language instructions. SOPs are specific to particular user languages rather than a 20 particular embodiment of CS 101, while conventional 't machine language instructions are specific to particular S, machines. SOPs are in turn interpreted and executed by t microcode. There will be an S-Language Dialect, a set of SOPs, for each user languages. CS 101, for example, may have SOP Dialects for COBOL, FORTRAN, and SPL. A particular distinction of CS 101 is that all SOPs are of a uniform, fixed length, for example 16 bits. CS 101 may J generally contain one or more sets of microcode for each i S-Language Dialect. These microcode Dialect Sets may be completely distinct, or may overlap where more than one SOP utilizes the same microcode.

As stated above, in CS 101 all operands are identified by Names, which are 8, 12, or 16 bit numbers.

CS 101 includes one or more "Name Tables" containing an Entry for each operand Name appearing in programs currently being executed. Each Name Table Entry contains information describing the operand referred to by a particular Name, and the directions necessary for CS 101 to translate that information into a corresponding logical descriptor. As previously described, logical descriptors may then be transformed into physical t 'descriptors to read and write operands from or to MEM 15 112. As described above, UIDs are unique for all CS 101 systems and AONs are unique within individual CS 101 systems. Names, however, are unique only within the context of a user's program. That is, a particular Name may appear 'in two different user's programs and, within 20 each program, will have different Name Table Entries and will refer to different operands.

,0 CS 101 may thereby be considered as utilizing two sets of instructions. A first set is comprised of SOPs, that is instructions selecting algorithms to be o l 25 executed. The second set of instructions are comprised of Names, which may be regarded as entry points into tables of instructions for making references regarding operands.

1 -41- Referring to Fig.3, a diagramic representation of CS 101 instruction stream is shown. A typical SIN is comprised of an SOP and may include one or more Names referring to operands. SOPs and Names allow user's programs to be expressed in very compact code. Fewer SOPs than machine language instructions are required to express a user's program. Also, use of SOPs allows easier and simpler construction of compilers, and facilitates adaption of CS 101 systems to new user languages. In addition, use of Names to refer to operands means that SOPs are independent of the form of the t, operands upon which they operate. This in turn allows for more compact code in expressing user programs in that SOPs specifying operations dependent upon operand form SS 15 are not required.

j 3. Archf-iectura--Base- r-terAdd igtq S1 As will be described further below, a user's program residing in CS 101 will include one or more Objects.

First, a Procedure Object contains at least the SINs of H 20 the user's programs and a Name Table containing entries Sl for operand Names of the program. The SINs may include references, or calls, to other Procedure Objects S- containing, for example, procedures available in common to many users. Second, a Static Data Area may contain static data, that is data having an existence for at least a single execution of the program. And third, a "i Macro-stack, described below, may contain local data, that is data generated during execution of a program.

Each Procedure Object, the Static Data Area and the Macro-stack are individual Objects identified by UIDs and i .r t: -42- AONs and addressable through UID and AON addresses and descriptors.

Locations of information within a user's Procedure Objects, Static Data Area, and Macro-stack are expressed as offsets from one of three values, or base addresses, referred to as Architectural Base Pointers (ABPs). For example, location information in Name Tables is expressed as offsets from one of the ABPs. ABPs may be expressed as previously described.

The three ABPs are the Frame Pointer the Procedure Base Pointer (PBP), and the Static Data Pointer (SDP). Locations of data local to a procedure, for example in the procedure's Macrostack, are described as offsets from FP. Locations of non-local data, that is Static Data, are described as offsets from SDP.

o Locations of SINs in Procedure Objects are expressed as offsets from PBP; these offsets are determined as a Program Counter (PC) value. Values of the ABPs vary during program execution and are therefore not provided 20 by the compiler converting a user's high level language program into a program to be executed in a CS 101 system. When the program is executed, CS 101 provides the proper values for the ABPs. When a program 'is actually being exected, the ABP's values are stored in FU 120's GRF.

-42an adrsal thog I n O drse n descriptors Locaionsof nforatin wihina usr'sProcdur -43- Other pointers are used, for example, to identify the top frame of CS 101's Secure Stack (a microcode level stack described below) or to identify the microcode Dialect currently being used in execute the SINs of a procedure. These pointers are similar to FP, SDP, and

PBP.

4. Stack Referring to Fig. 4 and 4A, diagramic representations of various control levels and stack mechanisms of, respectively, conventional machines and CS 101, are shown. Referring first to Fig. 4, top level of control is provided by User Language Instructions 402, for r example in FORTRAN or COBOL. User Language instructions S402 are converted into a greater number of more detailed Machine Language Instructions 404, used within a machine to execute user's programs. Within the machine, Machine Language Instructions 404 are interpreted and executed by Microcode Instructions 406, that is sequences of microinstructions which in turn directly control Machine B 20 Hardware 408. Some conventional machines may include a K Stack Mechanism 410 used to save current machine state, that is current microinstruction and contents of various

S

t machine registers, if a current Machine Language SInstruction 404 cannot be executed or is interrupted. In general, machine state on the microcode and hardware level is not saved. Execution of a current Machine m, Language Instruction 404 is later resumed at start of the microinstruction sequence for executing that Machine Language Instruction 404.

I

4 -44rIt I

C

t 44 4 ii 4l 4 4. ii ct 's C IC I 4. i.

t~ It Referring to Fig. 4A, top level control in CS 101 is by User Language Instructions 412 as in a conventional machine. In CS 101, however, User Language Instructions 412 are translated into SOPs 414 which are of a higher level than conventional machine language instructions.

In general, a single User Language Instruction 412 is transformed into at most two or three SOPs 414, as opposed to an entire sequence of conventional Machine Language Instructions 404. SOPs 414 are interpreted and executed by Microcode Instructions 416 (sequences of microinstructions) which directly control CS 101 Hardware 418. CS 101 includes a Macro-stack Mechanism (MAS) 420, at SOPs 414 level, which is comparable to but different in construction and operation from a conventional Machine Language Stack Mechanism 410. CS 101 also includes Micro-code Stack Mechanisms 422 operating at Microcode 416 level, so that execution of an interrupted microinstruction of a microinstruction sequence may be later resumed with the particular microinstruction which 20 was active at the time of the interrupt. CS 101 is therefore more efficient in handling interrupts in that execution of microinstruction sequences is resumed from the particular point that a microinstruction sequence was interrupted, rather from the beginning of that sequence.

As will be described further below, CS 101's Micro-code Stack Mechanisms 422 on microcode level is effectively comprised of two stack mechanisms. The first stack is Micro-instruction Stack (MIS) 424 while the second stack is referred to as Monitor Stack (MOS) 426. CS 101 SIN Microcode 428 and MIS 424 are primarily concerned with I1 execution of SOPs of user's programs. Monitor Microcode 430 and MOS 426 are concerned with operation of certain CS 101 internal functions.

Division of CS 101's microcode stacks into an MIS 424 and a MOS 426 illustrates a further feature of CS 101.

In conventional machines, monitor functions may be performed by a separate CPU operating in conjunction with the machine's primary CPU. In CS 101, a single hardware CPU is used to perform both functions with actual execution of both functions performed by separate groups of microcode. Monitor microcode operations may be initiated either by certain SINs 414 or by control t signals generated directly by CS 101's Hardware 418.

Invocation of Monitor Microcode 430 by Hardware 418 S 15 generated signals insures that CS 101's monitor functions may always be invoked.

Referring to Fig. 5, a diagramic representation of CS 101's stack mechanisms for a sin~ge use r--s rogam, or Srocedure 'is shown. Basically, and with exception of j 20 MOS 426, CS 101's stacks reside in MEM 112 with certain portions of those stacks accelerated into FU 120 and EU 122 to enhance speed of operation.

Certain areas of MEM 112 storage space are set aside to contain Macro-Stacks (MASs) 502, stack mechanisms operating on the SINs level, as described above. Other areas of MEM 112 are set aside to contain Secure Stack (SS) 504, operating on the microcode level, as described above and of which MIS 424 is a part.

i t -46- As described further below, both FU 120 and EU 122 contain register file arrays, referred to respectively as GRF and ERF, in addition to registers associated with, for example, ALUs. Referring to FU 120, shown therein is FU 120's GRF 506. GRF 506 is horizontally divided into three areas. A first area, referred to as General Registers (GRs) 508 may in general be used in the same manner as registers in a conventional machine. A second area of GRF 506 is Micro-Stack (MIS) 424, and is set aside to contain a portion of a Process's SS 504. A third portion of GRF 506 is set aside to contain MOS 426. Also indicated in FU 120 is a block referred to as Microcode Control State (mCS) 510. mCS 510 represents t* registers and other FU 120 hardware containing current 15 operating state of FU 120 on the microinstruction and i hardware level. mCS 510 may include, for example, the Scurrent microinstruction controlling operation of FU 120.

Referring to EU 122, indicated therein is a first block referred to as Execute Unit State (EUS) 512 and a second block referred to as SOP Stack 514. EUS 512 is similar to mCS 510 in FU 120 and includes all registers and other EU 122 hardware containing I, information reflecting EU 122's current operating state.

SOP Stack 518 is a portion of EU 122's ERF 516 which has been set aside as a stack mechanism to contain a portion of a process's SS 504 pertaining to EU 122 operations.

Considering first MASs 502, as stated above MASs 502 operate generally upon the SINs level. MASs 502 are used in general to store current state of a process's (defined below) execution of a user's program.

-47- Referring next to MIS 424, in a present embodiment of CS 101 that portion of GRF 506 set aside to contain MIS 424 may have a capacity of eight stack frames. That is, up to 8 microinstruction level interrupts or calls pertaining to execution of a user's program may be stacked within MIS 424. Information stored in MIS 424 stack frames is generally information from GR 508 and MCS 510. MIS 424 stack frames are transferred between MIS 424 and SS 504 such that at least one frame, and no more than 8 frames, of SS 504 reside in GRF 506. This insures that at least the top-most frames of a process's SS 504 are present in FU 120, thereby enhancing speed of operation of FU 120 by providing rapid access to those top frames. SS 504, residing in MEM 112, may contain, 15 for all practical purposes, an unlimited number of frames so that MIS 424 and SS 504 appear to a user to be .effectively an infinitely deep stack.

MOS 426 resides entirely in FU 120 and, in a present embodiment 'of CS 101, may have a capacity of 8 stack 20 frames. A feature of CS 101 operation is that CS 101 mechanisms for handling certain events or interrupts should not rely in its operation upon those portions of CS 101 whose operation has resulted in those faults or interrupts. Among events handled by CS 101 monitor microcode, for example, are MEM 112 page faults. An MEM 112 page fault occurs whenever FU 120 makes a reference to data in MEM 112 and that data is not in MEM 112. Due t i -48to this and similar operations, MOS 426 resides entirely in FU 120 and thus does not rely upon information in MEM 112.

As described above, GRs 508, MIS 424, and MOS 426 each reside in certain assigned portions of GRF 506.

This allows flexibility in modifying the capacity of GRs 508, MIS 424, and MOS 426 as indicated by experience, or to modify an individual CS 101 for particular purposes.

Referring finally to EU 122, EUS 512 is functionally a part of a process's SS 504. Also as previously described, EU 122 performs arithmetic operations in response to SINs and may be interrupted by FU 120 to aid certain FU 120 operations. EUS 512 allows stacking of interrupts. For example, FU 120 may first interrupt an 15 arithmetic SOP to request EU 122 to aid in evaluation of a Name Table Entry. Before that first interrupt is S. completed, FU 120 may interrupt again, and so on.

t SOP Stack 514, is a single frame stack for storing current sta'te of EU 122 when an interrupt interrupts 20 execution of an arithmetic SOP. An interrupted SOP's state is transferred into SOP Stack 514 and the interrupt begins execution in EUS 512. Upon occurrence of a second interrupt (before the first interrupt is completed) EU's first interrupt state is transferred from EUS 512 to a stack frame in SS 504, and execution of the second interrupt begins in EUS 512. If a third interrupt occurs before completion of second interrupt, EU's second interrupt state is transferred from EUS 512 to another stack frame in SS 504 and execution of the third interrupt is begun in EUS 512; and so on. EUS 512 and SS 504 thus provide an apparently infinitely deep microstack Si -49for EU 122. Assuming that the third interrupt is completed, state of second interrupt is transferred from SS 504 to EUS 512 and execution of second interrupt resumed. Upon completion of second interrupt, state of first interrupt is transferred from SS 504 to EUS 512 and completed. After completion of first interrupt, state of the original SOP is transferred from SOP Stack 514 to EUS 512 and execution of that SOP resumed.

C. Procedure--Processes-,--and-Vrttual--Pro 'ors Eia:-'€ Referring to Fig. 6, a diagramic representation of procedures, processes, and virtual processes is shown.

As described above, a user's program to be executed is compiled to result in a Procedure 602. A Procedure 602 S. 15 includes a User's Procedure Object 604 containing the SOPs of the user's program and a Name Table containing Entries for operand Names of the user's program, and a Static Data Area 606. A Procedure 602 may also include other Procedure Objects 608, for example utility programs e 20 available in common to many users. In effect, a Procedure 602 contains the instructions (procedures) and data of a user's program.

I t A Process 610 includes, as described above, a Macro- SStack (MAS) 502 storing state of execution of a user's 25 Procedure 602 at the SOP level, and a Secure Stack (SS) 504 storing state of execution of a user's Procedure 602 at the microcode level. A Process 610 is associated with r a user's Procedure 602 through the ABPs described above and which are stored in the MAS 502 of the Process 610.

Similarly, the MAS 502 and SS 504 of a Process 610 are

I

r ri.P" dL i #r I 4 4* 414 I 4" 44 i ci I *44$ 4* *i *49, *c 49 .t 94 *i 4 4* I.r 4 4 I 4

I-

IC IC associated through non-architectural pointers, described above. A Process 602 is effectively a body of information linking the resources, hardware, microcode, and software, of CS 101 to a user's Procedure 602. In effect, a Process 610 makes the resources of CS 101 available to a user's Procedure 602 for executing of that Procedure 602. CS 101 is a multi-program machine capable of accommodating up to, for example, 128 Processes 610 concurrently. The number of Processes 610 which may be executed concurrently is determined by the number of Virtual Processors 612 of CS 101. There may be, for example, up to 16 Virtual Processors 612.

As indicated in Fig. 6, a Virtual Processor 612 is comprised of a Virtual Processor State Block (VPSB) 614 associated with the SS 504 of a Process 612. A VPSB 614 is, in effect, a body of information accessible to CS 101's operating system and through which CS 101's operating system is informed of, and provided with access to, a Process 610 through that Process 610's SS 504. A 20 VPSB 614 is associated with a particular Process 610 by writing information regarding that Process 610 into that VPSB 614. CS 101's operating system may, by gaining access to a Process 610 through an associated PSB 614, read information, such as ABP's, from that Process 610 to FU 120, thereby swapping that Process 610 onto FU 120 for execution. It is said that a Virtual Processor 612 thereby executes a Process 610; a Virtual Processor 612 may be regarded therefor, as a processor having "Virtual", or potential, existence which becomes "real" 4 4 I r' fi-~ I- ~P

_I

I

3 r -51tt t 4 tr I t f* A.

U.

7 t1 when its associated Process 610 is swapped into FU 120.

In CS 101, as indicated in Fig. 6, only one Virtual Processor 612 may execute on FU 120 at a time and the operating system selects which Virtual Processor 612 will excecute on FU 120 at any given time. In addition, CS 101's operating system selects which Processes 610 will be associated with the available Virtual Processors 612.

Having briefly described certain individual structural and operating features of CS 101, the overall operation of CS 101 will be described in further detail next below in terms of these individual features.

D. CS 1 01--Overa-- Structutr--and-Operaton-f-Pcs---7-.

8, 9"i .10i--"tti."^21 3 1. .n.t roduction- Fiag'-7 4 As indicated in Fig. 7, CS 101 is a multiple level system wherein operations in one level are generally transparent to higher levels. User 701 does not see the S-Language, addressing, and protection mechanisms defined at Architectural Level 708. Instead, he sees User 20 Interface 709, which is defined by Compilers 702, Binder 703, and Extended (high level) Operating System (EOS) 704. Compilers 702 translate high-level language code into SINs and Binder 703 translates symbolic Names in programs into UID-offset addresses.

As Fig. 7 shows, Architectural Level 708 is not defined by FU 120 Interface 711. Instead, the architectural resources level are created by S-Language interpreted SINs when a program is executed; Name f -52- Interpreter 715 operates under control of S-Language Interpreters 705 and translates Names into logical descriptors. In CS 101, both S-Language Interpreters 705 and Name Interpreter 715 are implemented as microcode which executes on FU 120. S-Language Interpreters 705 may also use EU 122 to perform calculations. A Kernel Operating System (KOS) provides CS 101 with UID-offset addressing, objects, access checking, processes, and virtual processors, described further below. KOS has three kinds of components: KOS Microcode 710, KOS Software 706, and KOS Tables in MEM 112. KOS 710 components are microcode routines which assist FU 120 in performing certain required operations. Like other high- *level language routines, KOS 706 components contain SINs jj* *15 which are interpreted by S-Interpreter Microcode 705.

S* Many KOS High-Level Language Routines 706 are executed by C IA special KOS processes; others may be executed by any process. Both KOS High-Level Language Routines 706 and KOS Microcode 710 manipulate KOS Tables in MEM 112.

S- FU 120 Interface 711 is visible only to KOS and to S- Interpreter Microcode 705. For the purposes of this t* discussion, FU 120 may be seen as a processor which contains the following main elements: A Control Mechanism 725 which executes 2' 5 microcode stored in Writable Control Store 713 and manipulates FU 120 devices as directed by this microcode.

-53- A GRF 506 containing registers in which data may be stored.

A Processing Unit 715.

All microcode which executes on FU 120 uses these devices; there is in addition. a group of devices for performing special functions; these devices are used only by microcode connected with those functions. The microcode, the specialized devices, and sometimes tables in MEM 112 make up logical machines for performing certain functions. These machines will be described in detail below.

In the following, each of the levels illustrated in us., Fig. 7 will be discussed in turn. First, the components at User Interface 709 will be examined to see how they 15 translate user programs and requests into forms usable by j CS 101. Then the components below the User Interface 709 will be examined to sea how they create logical machines for performing CS 101 operations.

2. Compmilers 70-t2,- 7.

Compilers 702 translate files containing the highlevel language code written by User 701 into Procedure Objects 608. Two components of a Procedure Object 608 are code (SOPs) and Names, previously described. SOPs represent operations, and the Names represent data. A single SIN thus specifies an operation to be performed on the data represented by the Names.

e C t, -54- 3. Binder- In some cases, Compiler 702 cannot define locations as offsets from an ABP. For example, if a procedure calls a procedure contained in another procedure object, the location to which the call transfers control cannot be defined as an offset from the PBP used by the calling procedure. In these cases, the compiler uses symbolic Names to define the locations. Binder 703 is a utility which trar;alates symbolic Names into UID-offset addresses. It does so in two ways: by combining separate Procedure Objects 608 into a single large Procedure Object 608, and then redefining symbolic Names as offsets from that Procedure Object 608's ABPs, or by

S

e translating symbolic Names when the program is executed.

In the second case, Binder 703 requires assistance from EOS 704.

4. EOS "704 -f Pia-7-) EOS 704 manages the resources that User 701 requires to execute his programs. From User 701's point of view, 20 the most important of these resources are files and I processes. EOS 704 creates files by requesting KOS to create an object and then mapping the file onto the K ,Kc object. When a User 701 performs an operation on a file, EOS 704 translates the file operation into an operation on an object. KOS creates them at EOS 704's request and makes them available to EOS 704, which in turn makes them available to User 701. EOS 704 causes a process .e execute by associating it a Virtual Processor 612. In >4 1 v 1 logical terms, a Virtual Processor 612 is the means which KOS provides EOS 704 for executing Processes 610. As many Processes 610 may apparently execute simultaneously in CS 101 as there are Virtual Processors 612. The illusion of simultaneous execution is created by multiplexing JP 114 among the Virtual Processors; the manner in which Processes 610 and Virtual Processors 610 are implemented will be explained in detail below.

KOS'-and--Architectura-l--err1 face-7 8 <Piqr-.

S-Interpreter Microcode 710 and Name Interpreter Microcode 715 require an environment provided by .KOS Microcode 710 and KOS Software 706 to execute SINs. For example, as previously explained, Names and program 15 locations are defined in terms of ABPs whose values vary during execution of the program. The KOS environment provides valuea for the ABPs, and therefore makes it possible to interpret Names and program locations as locations in MEM 112. Similarly, KOS help is required to transform logical descriptors into references to MEM 112 and to perform protection checks.

The environment provided by KOS has the following elements: 'e A Process 610 which contains the state of an execution of the program for a given User 701.

A Virtual Processor 612 which gives the Process 610 access to JP 114.

r: II4 r

I

-56- 9tte 9 9* 99 9 9 99 9I 9 9e 9 994

S'

I

A 94 9 4* *4 *999 4 .4 94 4 9 1* An Object Management System which translates UIDs into values that are usable inside JP 114.

A Protection System which checks whether a Process 610 has the right to perform an operation on an Object.

A Virtual Memory Management System which moves those portions of Objects which a Process 610 actually references from the outside world into MEM 112 and translates logical descriptors into physical descriptors.

In the following, the logical properties of this environment and the manner in which a program is executed 15 in it will be explained.

6. Procsesl ea-6--and-Vi7ta -Proce-osgrs--l-2 Processes 610 and Virtual Processors 612 have already been described in logical terms; Fig. 8 gives a highlevel view of their physical implementation.

Fig. 8 illustrates the relationship between Processes 610, Virtual Processors 612, and JP 114. In physical terms, a Process 610 is an area of MEM 112 which contains the current state of a user's execution of a program.

25 One example of such state is the current values of the ABPs and a Program Counter Given the current value of the PBP and the PC, the next SOP in the program can be executed; similarly, given the current values of SDP and y,£tI

A

I

c -57- FP, the program's Names can be correctly resolved. Since the Process 610 contains the current state of a program's execution, the program's physical execution can be stopped and resumed at any point. It is thus possible to control program execution by means of the Process 610.

As already mentioned, a Process 610's execution proceeds only when KOS has bound it to a Virtual Processor 612, that is, an area of MEM 112 containing the state required to execute microinstructions on JP 114 hardware. The operation of binding is simply a transfer of Process 610 state from the Process 610's area of MEM 112 to a Virtual Processor 612's area of MEM 112. Since binding and unbinding may take place at any time, EOS 704 may multiplex Processes 610 among Virtual Processors 612. In Fig. 8, there are more Processes 610 than there 4 are Virtual Processors 612. The physical execution of a Process 610 on JP 114 takes place only while the Process 610's Virtual Processor 612 is bound to JP 114, i.e., when state is transferred from Virtual Processor 612's area of MEM 112 to JP 114's registers. Just as EOS 704 multiplexes Virtual Processors 612 among Processes 610, KOS multiplexes JP 114 among Virtual Processors 612. In Fig. 8, only one Process 610 is being physically

S**

1 executed. The means by which JP 114 is multiplexed among Virtual Processors 612 will be described in further detail below.

vl -58- 7. Processes 610--and-Stacke FPia.9 In CS 101 systems, a Process 610 is made up of six Objects: one Process Object 901 and Five Stack Objects 902 to 906. Fig. 9 illustrates a Process 610. Process Object 901 contains the information which EOS 704 requires to manage the Process 610. EOS 704 has no direct access to Process Object 901, but instead obtains the information it needs by means of functions provided to it by KOS 706, 710. Included in the information are the UIDs of Stack Objects 902 through 906. Stack Objects 902 to 906 contain the Process 610's state.

Stack Objects 902 through 905, are required by CS 101's domain protection method and comprise Process 610's MAS 502. Briefly, a domain is determined in part by t 15 operations performed when a system is operating in that 'I domain. For example, the system is in EOS 704 domain when executing EOS 704 operations and in KOS 706, 710 domain when executing KOS 706, 710 operations. A Process 510 must have one stack for each domain it enters. In the present embodiment, the number of domains is fixed at *i four, but alternate embodiments may allow any number of *domains, and correspondingly, any number of Stack Objects. Stack Object 906 comprises Process 610's Secure i Stack 504 and is required to store state which may be manipulated only by KOS 706, 710.

Each invocation made by a Process 610 results in the

SC

l addition of frames to Secure Stack 504 and to Macro-Stack te 502. The state stored in the Secure Stack 504 frame I -59includes the macrostate for the invocation, the state required to bind Process 610 to a Virtual Processor 612.

The frame added to Macro-Stack 502 is placed in one of Stack Objects 902 through 905. Which Stack Objects 902 to 905 gets the frame is determined by the invoked procedure's domain of execution.

Fig. 9 shows the condition of a Process 610's MAS 502 and Secure Stack 504 after the Process 610 has executed four invocations. Secure Stack 504 has one frame for each invocation; the frames of Process 610's MAS 502 are found in Stack Objects 902, 904, and 905. As revealed by their locations, Frame 1 is for an invocation of a routine with KOS 706, 710 domain of execution, Frame 2 for an invocation of a routine with the EOS 704 domain of -15 execution, and Frames 3 and 4 for invocations of routines with the User domain of execution. Process 610 has not yet invoked a routine with the Data Base Management 4 System (DBMS) domain of execution. The frames in Stack Objects 902 through 905 are linked together, and a frame is added to or removed from Secure Stack 504 every time a frame is added to Stack Objects 902 through 905. MAS 502 and Secure Stack 504 thereby function as a single logical stack even though logically contained in five separate 4 Objects.

8. Processe s-6-O--aend-a^.-Tlf In the CS 101, calls and returns are executed by KOS 706, 710. When KOS 706, 710 performs a call for a process, it does the following: 1I t w 4 S t 4 t I i 1 4 1# 4 4 (4 It saves the calling invocation's macrostate in the top frame of Secure Stack 504 (Fig. 9).

It locates the procedure whose Name is contained in the call. The location of the first SIN in the procedure becomes the new

PBP.

Using information contained in the called procedure, KOS 706, 710 creates.a new MAS 502 frame in the proper Stack Object 902 through 905 and a new Secure Stack 504 frame in Secure Stack 504. FP is updated to point to the new MAS 502. If necessary, SDP is also updated.

Once the values of the ABPs have been updated, the PC is defined, Names can be resolved, and execution of the invoked routine can commence. On a return from the invocation to the invoking routine, the stack frames are deleted and the ABPs are set to the values saved in the invoking routine's macrostate. The invoking routine then continues execution at the point following the invocation.

A Process 610 may be illustrated in detail by putting the FORTRAN statement A B into a FORTRAN routine called EXAMPLE and invoking it from another FORTRAN routine named CALLER. To simplify the example, it is assumed that CALLER and EXAMPLE both have the same domain of

L

ilri i I- i .1 4 I -61execution. The parts of EXAMPLE which are of interest look like this: SUBROUTINE EXAMPLE (C) INTEGER X,C INTEGER A,B A =B

RETURN

END

The new elements are a formal argument, C, and a new local variable, X. A formal argument is a data item which receives its value from a data item used in the invoking routine. The formal argument's value thus t* 15 varies from invocation to invocation. The portions of S, INVOKER which are of interest look like this: SUBROUTINE INVOKER 444$40 INTEGER Z .4 CALL EXAMPLE (Z) *444

END

4 The CALL statement in INVOKER specifies the Name of i the subroutine being invoked and the actual arguments for the subroutine's formal arguments. During the invocation, the subroutine's formal arguments take on the ,rr values of the actual arguments. Thus, during the invocation specified by this CALL statement, the formal 4 -62argument C will have the value represented by the variable Z in INVOKER.

When INVOKER is compiled, the compiler produces a CALL SIN corresponding to the CALL statement. The CALL SIN contains a Name representing a pointer to the beginning of the called routine's location in a procedure object and a list of Names representing the call's actual arguments. When CALL is executed, the Names are interpreted to resolve the SIN's Names as previously described, and KOS 710 microcode to perform MAS 502 and Secure Stack 504 operations.

Fig. 10 illustrates the manner in which the KOS 710 call microcode manipulates MAS 502 and Secure Stack 504.

Fig. 10 includes the following elements: Call Microcode 1001, contained in FU 120 Writable Control Store 1014.

PC Device 1002, which contains part of macrostate belonging to the invocation of INVOKER which is executing the CALL statement.

Registers in FU Registers 1014. Registers 1004 contents include the remainder of macrostate and the descriptors S t corresponding to Names for EXAMPLE's location and the actual argument Z.

Procedure Object 1006 contains the entries C for INVOKER and EXAMPLE, their Name Tables, Ir,; and their code.

i -63- Macro-Stack Object 1008 (MAS 502) and Secure Stack Object 1010 (Secure Stack 504) contain the stack frames for the invocations of INVOKER and EXAMPLE being discussed here. EXAMPLE's frame is in the same Macro-Stack object as INVOKER's frame because both routines are contained in the same Procedure Object 1006, and therefore have the same domain of execution.

KOS Call Microcode 1001 first saves the macrostate of INVOKER's invocation on Secure Stack 504. As will be discussed later, when the state is saved, KOS 706 Call Microcode 1001 uses other KOS 706 microcode to translate T: the location information contained in the macrostate into 4' 15 the kind of pointers used in MEM 112. Then Microcode 1001 uses the descriptor for the routine Name to locate the pointer to EXAMPLE's entry in Procedure Object 1006.

S" From the entry, it locates pointers to EXAMPLE's Name Table and the beginning of EXAMPLE's code. Microcode 1001 takes these pointers, uses other KOS 706 microcode i. to translate them into descriptors, and places the descriptors in the locations in Registers 1004 reserved for the values of the PBP and NTP. It then updates the t' values contained in PC Device 1002 so that when the call is finished, the next SIN to be executed will be the first SIN in EXAMPLE.

'1 4 I i p- -64- CALL Microcode 1001 next constructs the frames for EXAMPLE on Secure Stack 504 and Macro-Stack 502. This discussion concerns itself only with Frame 1102 on Macro- Stack 502. Fig. 11 illustrates EXAMPLE's Frame 1102.

The size of Frame 1102 is determined by EXAMPLE's local variables A, and B) and formal arguments At the bottom of Frame 1102 is Header 1104. Header 1104 contains information used by KOS 706, 710 to manage the stack. Next comes Pointer 1106 to the location which contains the value represented by the argument C. In the invocation, the actual for C is the local variable Z in INVOKER. As is the case with all local variables, the storage represented by Z is contained in the stack frame belonging to INVOKER's invocation. When a name S 15 interpreter resolved C's name, it placed the descriptor in a register. Call Microcode 1001 takes this descriptor, converts it to a pointer, and stores the pointer above Header 1104.

Since the FP ABP points to the location following the last pointer to an actual argument, Call Microcode 1001 can now calculate that location, convert it into a descriptor, and place it in a FU Register 1004 reserved for VP. The next step is providing storage for EXAMPLE's 4" local variables. EXAMPLE'S Procedure Object 1006 contains the size of the storage required for the local variables, so Call Microcode 1001 obtains this information from Procedure object 1006 and adds that much storage to Frame 1102. Using the new value of FP and the information contained in the Name Table Entries for the local data, Name Interpreter 715 can now construct descriptors for the local data. For example, A's entry in Name Table specified that it was offset 32 bits from FP, and was 32 bits long. Thus, its storage falls between the storage for X and B in Figure 11.

9. Memor yReferences--td-the Virtual M-emoy Management--Satem -12 I- As already explained, a logical descriptor contains an AON field, an offset field, and a length field. Fig.

12 illustrates a Physical Descriptor. Physical Descriptor 1202 contains a Frame Number (FN) field, a Displacement field, and a Length field.

Together, the Frame Number field and the Displacement t 15 field specify the location in MEM 112 containing the data, and the Length field specifies the length of the data.

As is clear from the above, the virtual memory management system must translate the AON-offset location contained in a logical descriptor 1204 into a Frame Number-Displacement location. It does so by associating t logical pages with MEM 112 frames. MEM 112 frames are not to be confused with stack frames). Fig. 13, S|illustrates how Macrostack 502 Object 1302 is divided into Logical Pages 1304 in secondary memory and how Logical Pages 1304 are moved onto Frames 1306 in MEM 112. A Frame 1306 is a fixed-size, contiguous area of MEM 112. When the virtual memory management system -66brings data into MEM 112, it does so in frame-sized chunks called Logical Pages 1308. Thus, from the virtual memory system's point of view, each object is divided into Logical Pages 1308 and the address of data on a page consists of the AON of the data's Object, the number of pages in the object, and its displacement on the page.

In Fig. 13, the location of the local variable B of EXAMPLE is shown as it is defined by the Virtual memory system. B's location is a UID and an offset, or, inside JP 114, an AON and an offset. As defined by the virtual memory system, B's location is the AON, the page number 1308, and a displacement within the page. When a process references the variable B, the virtual memory management t system moves all of Logical Page 1308 into a MEM 112 Frame 1306. B's displacement remains the same, and the i virtual memory system translates its Logical Page Number 1308 into the number of Frame 1306 in MEM 112 which S* contains the page.

The virtual memory management system must therefore perform two kinds of translations: AON-offset I addresse into AON-plge number-displacement addresses, f t and AON-pgaq te inLto a frame number.

Access-ont o-f-Piq-r-+1 t Each time a reference is made to an Object, KOS 706, 710 checks whether the reference is legal. The following discusson will first present the logical structure of access control in CS 101, and then discuss the microcode and devices which implement it.

-67- CS 101 defines access in terms of subjects, modes of access, and Object size. A process may reference a data item located in an Object if three conditions hold: 1) If the process's subject has access to the Object.

2) If the modes of access specified for the subject include those required to perform the intended operation.

3) If the data item is completely contained in the Object, if the data item's length added to the data item's offset do not exceed the number of bits in the Object.

The subjects which have access to an Object and the ,itti kinds of access they have to the Object are specified by a data structure associated with the Object called the r* Access Control List (ACL). An Object's size is one of its attributes. Neither an Object's size nor its ACL is i contained in the Object. Both are contained in system tables, and are accessible by means of the Object's UID.

Fig. 14 shows the logical structure of access control in CS 101. Subject 1408 has four components: Principal J 1404, Process 1405, Domain 1406, and Tag 1407. Tag 1407 is not implemented in a present embodiment of CS 101, so V t the following' description will deal only with Principal 1404, Process 1405, and Domain 1406.

Principal 1404 specifies a user for which the process which is making the reference was created; -68- Process 1405 specifies the process which is making the reference; and, Domain 1406 specifies the domain of execution of the procedure which the process is executing when it makes the reference.

Each component of the Subject 1408 is represented by a UID. If the UID is a null UID, that component of the subject does not affect access checking. Non-null UIDs are the UIDs of Objects that contain information about the subject components. Principal Object 1404 contains identification and accounting information regarding system users, Process Object 1405 contains process management information, and Domain Object 1406 contains information about per-domain error handlers.

I There may be three modes of accessing an Object 1410: S ,tit. read, write, and execute. Read and write are selfexplanatory; execute is access which allows a subject to execute instructions contained in the Object.

Access Control Lists (ACLs) 1412 are made up of U Entries 1414. Each entry two components: Subject Template 1416 and Mode Specifier 1418. Subject Template 4 1416 specifies a group of subjects that may reference the S* Object and Mode Specifier 1418 specifies the kinds of access these subjects may have to the Object. Logically speaking, ACL 1412 is checked each time a process references an Object 1410. The reference may succeed I only if the process's current Subject 1408 is one of It i4 -69those on Object 1410's ACL 1412 and if the modes in the ACL Entry 1414 for the Subject 1408 allow the kind of access the process wishes to make.

11. V-itua-ti-Procssaeta-and- Virtu-rcal ee Swappting-I--f As previously mentioned, the execution of a program by a Process 610 cannot take place unless EOS 704 has bound the Process 610 to a Virtual Processor 612.

Physical execution of the Process 610 takes place only while the process's Virtual Processor 612 is bound to JP 114. The following discussion deals with the data bases belonging to a Virtual Processor 612 and the means by which a Virtual Processor 612 is bound to and removed from JP 114.

15 Fig. 15 illustrates the devices and tables which KOS *706, 710 uses to implement Virtual Processors 612. FU 120 WCS contains KOS Microcode 706 for binding Virtual SProcessors 612 to JP 114 and removing them from JP 114.

Timers 1502 and Interrupt Line 1504 are hardware devices which produce signals that cause the invocation of KOS .Microcode 706. Timers 1502 contains two timing devices: Interval Timer 1506, which may be set by KOS 706, 710 to signal when a certain time is reached, and Egg Timer S 1508, which guarantees that there is a maximum time interval for which a Virtual Processor 612 can be bound to JP 114 before it invokes KOS Microcode 706. Interrupt 1 Line 1504 becomes active when JP 114 receives a message i from IOS 116, for example when IOS 116 has finished loading a logical page into MEM 112.

c j v FU 120 Registers 508 contain state belonging to the Virtual Processor 612 currently bound to JP 114. Here, this Virtual Processor 612 is called Virtua'. Processor A. In addition, Registers 508 contain registers reserved for the execution of VP Swapping Microcode 1510. ALU 1942 (part of FU 120) is used for the descriptor-topointer and pointer-to-descriptor transformations required when one Virtual Processor 612.is unbound from JP 114 and another bound to JP 114. MEM 112 contains data bases for Virtual Processors 612 and data bases used by KOS 706, 710 to manage Virtual Processors 612. KOS 706, 710 provides a fixed number of Virtual Processors 612 for CS 101. Each Virtual Processor 612 is represented by a Virtual Processor State Block (VPSB) 614. Each VPSB .5 614 contains information used by KOS 706, 710 to manage f t the Virtual Processor 612, and in addition contains information associating the Virtual Processor 612 with a process. Fig. 15 shows two VPSBs 614, one belonging to Virtual Processor 612A, and another belonging to Virtual Processor 612B, which will replace Virtual Processor 612A on JP 114. The VPSBs 614 are contained in VPSB Array 1512. The index of a VPSB 614 in VPSB Array 1512 is Virtual Processor Number 1514 belonging to the Virtual t Processor 612 represented by a VPSB 614. Virtual Processor Lists 1516 are lists which KOS 706, 710 uses to manage Virtual Processors 612. If a Virtual Processor 612 is able to execute, its Virtual Processor Number 1514 is on a list called the Runnable List; Virtual Processors 1- j J r, *i -71- 612 which cannot run are on other lists, depending on the reason why they cannot run. It is assumed that Virtual Processor 612B's Virtual Processor Number 1514 is the first one on the Runnable List.

When a process is bound to a Virtual Procesor 612, the Virtual Processor Number 1514 is copied into the process's Process Object 901 and the AONs of the process's Process Object 901 and stacks are copied into the Virtual Processor 612's VPSB 614. (AONs are used because a process's stacks are wired active as long as the process is bound to a Virtual Processor 612).

Binding is carried out by KOS 706, 710 at the request of EOS 704. In Fig. 15, two Secure Stack Objects 906 are .ctt shown, one belonging to the process to which Virtual Processor 612A is bound, and one belonging to that to which Virtual Processor 612B is bound.

SHaving described certain overall operating features of CS 101, a present implementation of CS 101's structure will be described further next below.

E. CS'.1 .:-Stuctua^m---I m mentation Fia-.

Referring to Fig. 16, a partial block diagram of IOS 116 is shown. Major elements of IOS 116 include an ECLIPSE* Burst Multiplexer Channel (BMC) 1614 and a NOVA' Data Channel (NDC) 1616, an 10 Controller (IOC) 1618 and a Data Mover (DM) 1610. IOS 116's data channel devices, for example BMC 1614 and NDC 1616, comprise IOS l16's interface to the outside world. Information and addresses are received from external devices, such as I I -72disk drives, communications modes, or other computer systems, by IOS 116's data channel devices and are transferred to DM 1610 (described below) to be written into MEM 112. Similarly, information read from MEM 112 is provided through DM 1610 to IOS 116's data channel devices and thus to the above described external devices. These external devices are a part of CS 101's addressable memory space and may be addressed through UID addresses.

IOC 1618 is a general purpose CPU, for example an ECLIPSE* computer available from Data General Corporation. A primary function of IOC 1618 is control of data transfer through IOS 116. In addition, IOC 1618 generates individual Maps for each data channel device :se 15 for translating external device addresses into physical addresses within MEM 112. As indicated in Fig. 16, each t~t data channel device contains an individual Address Translation Map (MAP) 1632 and 1636. This allows IOS 116 to assign individual areas of MEM 112's physical address space to each data channel device. This feature provides protection against one data channel device writing into or reading from information belonging to another data channel device. In addition, IOC 1618 may generate overlapping address translation Maps for two or more data 1 channel devices to allow these data channel devices to share a common area of MEM 112 physical address space.

L I

L

lO -73- Data transfer between IOS 116's data channel devices and MEM 112 is through DM 1610, which includes a Buffer memory (BUF) 1641. BUF 1641 allows MEM 112 and IOS 116 to operate asychronously. DM 1610 also includes a Ring Grant Generator (RGG) 1644 which controls access of various data channel devices to MEM 112. RGG 1644 is designed to be flexible in apportioning access to MEM 112 among IOS 116's data channel devices as loads carried by various data channel devices varies. In addition, RGG 1644 insures that no one, or group, of data channel devices may monopolize access to MEM 112.

Referring to Fig. 17, a diagramic representation of RGG 1644's operation is shown. As described further in a following description, RGG 1644 may be regarded as a commutator scanning a number of ports which are assigned t to various channel devices. For example, ports A, C, E, f and G may be assigned to a BMC 1614, ports B and F to a S• NDC 1616, and ports D and H to another data channel device. RGG 1644 will scan each of these ports in turn and, if the data channel device associated with a S* particular port is requesting access to MEM 112, will grant'access to MEM 112 to that data channel device. If no request is present at a given port, RGG 1644 will 1 I 6 continue immediately to the next port. Each data channel device assigned one or more ports is thereby insured opportunity of access to MEM 112. Unused ports, for j example indicating data channel devices which are not presently engaged in information transfer, are 1- -74effectively skipped over so that access to MEM 112 is dynamically modified according to the information transfer loads of the various data channel devices. RGG 1644's ports may be reassigned among IOS 116's various data channel devices as required to suit the needs of a particular CS 101 system. If, for example, a particular CS 101 utilizes NDC 1616 more than a BMC 1614, that CS 101's NDC 1616 may be assigned more ports while that CS 101's BMC 1614 is assigned fewer ports.

2. Memory--MEMf 2--Pi i1 Referring to Fig. 18, a partial block diagram of MEM 112 is shown. Major elements of MEM 112 are Main Store Bank (MSB) 1810, a Bank Controller (BC) 1814, a Memory Cache (MC) 1816, a Field Interface Unit (FIU) 1820, and Memory Interface Controller (MIC) 1822. Interconnections of these elements with input and output buses of MEM 112 to IOS 116 and JP 114 are indicated.

MEM 112 is an intelligent, prioritizing memory having a single port to IOS 116, comprised of IOM Bus 130, MIO Bus 129, and IOMC Bus 131, and dual ports to JP 114. A first JP 114 port is comprised of MOD Bus 140 and PD Bus 146, and a second port is comprised of JPD Bus 142 and PD Bus 146. In general, all data transfers from and to MEM t 112 by IOS 116 and JP 114 are of single, 32 bit words; IOM Bus 130, MIO Bus 129, MOD Bus 140, and JPD Bus 142 are each 32 bits wide. CS 101, however, is a variable I word length machine wherein the actual physical width of BI f data buses are not apparent to a user. For example, a t Name in a user's program may refer to an operand containing 97 bits of data. To the user, that 97 bit data item will appear to be read from MEM 112 to JP 114 in a single operation. In actuality, JP 114 will read that operand from MEM 112 in a series of read operations referred to as a string transfer. In this example, the string transfer will comprise three 32 bit read transfers and one single bit read transfer. The final single bit transfer, containing a single data bit, will be of a 32 bit word wherein one bit is data and 31 bits are fill.

Write operations to MEM 112 may be performed in the same manner. If a single read or write request to MEM 112 specifies a data item of less than 32 bits of data, that transfer will be accomplished in the same manner as the final transfer described above. That is, a single 32 bit word will be transferred wherein non-data bits are fill bits.

Bulk data storage in MEM 112 is provided in MSB 1810, which is comprised of one or more Memory Array cards (MAs) 1812. The data path into and out of MA 1812 is functions for MAs 1812. BC 1814's functions include addressing, transfer of data, controlling whether a read or write operation is performed, refresh, sniffing, and error correction code operations. All read and write operations from and to MAs 1812 through BC 1814 are in blocks of four 32 bit words.

S%,

1 I- l -76- The various MAs 1812 comprising MSB 1810 need not be of the same data storage capacity. For example, certain MAs 1812 may have a capacity of 256 kilobytes while other MAs 1812 may have a capacity of 512 kilobytes.

Addressing of the MAs 1812 in MSB 1810 is automatically adapted to various MA 1812 configurations. As indicated in Fig. 18, each MA 1812 contains an address circuit (A) which receives an input from the next lower MA 1812 indicating the highest address in that next lower MA 1812. The A circuit on an MA 1812 also receives an input from that MA 1812 indicating the total address space of that MA 1812. The A circuit of that MA 1812 adds the highest address input from next lower MA 1812 to its own ^input representing its own capacity and generates an output to the next MA 1812 indicating its own highest address. All MAs 1812 of MSB 1810 are addressed in r- parallel by BC 1814. Each MA 1812 compares such addresses to its input from the next lower MA 1812, representing highest address of that next lower MA 1812, and its own output, representing its own highest address, to determine whether a particular address provided by BC '1814 lies within the range of addresses contained within that particular MA 1812. The particular MA 1812 whose address space includes that address will then respond by 25 accepting the read or write request from BC 1814.

C C.

w -77- MC 1816 is the data path for transfer of data between BC 1814 and IOS 116 and JP 114. MC 1816 contains a high speed cache storing data from MSB 1810 which is currently being utilized by either IOS 116 or JP 114. MSB 1810 thereby provides MEM 112 with a large storage capacity while MC 1816 provides the appearance of a high speed memory. In addition to operating as a cache, MC 1816 includes a bypass write path which allows IOS 116 to write blocks of four 32 bit words directly into MSB 1810 through BC 1814. In addition, MC 1816 includes a cache write-back path which allows data to be transferred out of MC 1816's cache and stored while further data is transferred into MC 1816's cache. Displaced data from MC 1816's cache may then be written back into MSB 1810 at a 15 later, more convenient time. This write-back path enhances speed of operation of MC 1816 by avoiding delays incurred by transferring data from MC 1816 to MSB 1810 before new data may be written into MC 1816.

MEM 112's FIU 1820 allows manipulation of data formats in writes to and reads from MEM 112 by both JP S. 114 and IOS 116. For example, FIU 1820 may convert unpacked decimal data to packed decimal data, and vice S versa. In addition, FIU 1820 allows MEM 112 to operate as a bit addressable memory. For example, as described all data transfers to and from MEM 113 are of 32 bit words. If a data transfer of less than 32 bits is V required, the 32 bit word containing those data bits may

-I

4 A 7 -78- **4 4R 9 .9 6*9 4$ 9 S 4 4 *f be read from MC 1816 to FIU 1820 and therein manipulated to extract the required data bits. FIU 1820 then generates a 32 bit word containing those required data bits, plus fill bits, and provides that new 32 bit word to JP 114 or IOS 116. When writing into MEM 112 from IOS 116 through FIU 1820, data is transferred onto IOM Bus 130, read into FIU 1820, operated upon, transferred onto MOD Bus 140, and transferred from MOD Bus 140 to MC 1816. In read operations from MEM 112 to IOS 116, data is transferred from MC 1816 to MOD Bus 140, written into FIU 1820 and operated upon, and transferred onto MIO Bus 129 to IOS 116. In a data read from MEM 112 to JP 114, data is transferred from MC 1816 onto MOD Bus 140, transferred into FIU 1820 and operated upon, and 15 transferred again onto MOD Bus 140 to JP 114. In write operations from JP 114 to MEM 112, data on JPD Bus 142 is transferred into FIU 1820 and operated upon? and is then transferred onto MOD Bus 140 to MC 1816. MOD Bus 140 is thereby utilized as an MEM 112 internal bus for FIU 1820 20 operations.

Finally, MIC 1822 provides primary control of BC 1814, MC 1816, and FIU 1820. MIC 1822 receives control inputs from and provides control outputs to PD Bus 146 and IOMC Bus 131. MIC 1822 contains primary microcode control for MEM 112, but BC 1814, MC 1816, and FIU 1820 each include internal microcode control. Independent, internal microcode controls allow BC 1814, MC 1816, and FIU 1820 to operate independently of MIC 1822 after their

A

4

I-

-79operations have been initiated by MIC 1822. This allows BC 1814 and MSB 1810, MC 1816, and FIU 1820 to operate independently and asynchronously. Efficiency and speed of operation of MEM 112 are thereby enhanced by allowing pipelining of MEM 112 operations.

3. Petch-gUnit--+Fu 'ft-2t6f A primary function of FU 120 is to execute SINs. In doing so, FU 120 fetches instructions and data (SOPs and Names) from MEM 112, returns results of operations to MEM 112, directs operation of EU 122, executes instructions of user's programs, and performs the various functions of CS 101's operating systems. As part of these functions, FU 120 generates and manipulates logical addresses and descriptors and is capable of operating as a general purpose CPU.

S

f Referring to Fig. 19, a major element of FU 120 is the Descriptor Processor (DESP) 1910. DESP 1910 includes General Register File (GRF) 506. GRF 506 is a large register array divided vertically into three parts which are addressed in parallel. A first parte AONGRF 1932 stores AON fields of logical addresses and descriptors.

A second part, OFFGRF 1934, stores offset fields of logical addresses and descriptors and is utilized as a 32 Sbit wide general register array. A third portion GRF

S

v 25 506, LENGRF 1936, is a 32 bit wide register array for storing length fields of logical descriptors and as a general register for storing data. Primary data path from MEM 112 to FU 120 is through MOD Bus 140, which provides inputs to OFFGRF 1934. As indicated in Fig. 19, data may be transferred from OFFGRF 1934 to inputs of AONGRF 1932 and LENGRF 1936 through various interconnections. Similarly, outputs from LENGRF 1936 and AONGRF 1932 may be transferred to inputs of AONGRF 1932, OFFGRF 1934, and LENGRF 1936.

Output of OFFGRF 1934 is connected to inputs of DESP 1910's Arithmetic and Logic Unit (ALU) 1942. .ALU 1942 is a general purpose 32 bit ALU which may be used in generating and manipulating logical addresses and descriptors, as distinct from general purpose arithmetic and logic operands performed by MUX 1940. Output of ALU 1942 is connected to JPD Bus 142 to allow results of arithmetic and logic operations to be transferred to MEM 112 or EU 122.

t Also connected from output of OFFGRF 1934 is i*"r Descriptor Multiplexer (MUX) 1940. An output of MUX 1940 is provided to an input of ALU 1942. MUX 1940 is a 32 bit ALU, including an accumulator, for data manipulation operations. MUX 1940, together with ALU 1942, allows DESP 1910 to perform 32 bit arithmetic and logic felV* operations. MUX 1940 and ALU 1942 may allow arithmetic and logic operations upon operands of greater than 32 bits by performing successive operations upon successive 32 bit words of larger operands.

t, Logical descriptors or addresses generated or Sprovided by DESP 1910, are provided to Logical Descriptor zi.~ :-~13C -81a r ri a ,C a a.

*r a (LD) Bus 1902. LD Bus 1902 in turn is connected to an input of Address Translation Unit (ATU) 1928. ATU 1928 is a cache mechanism for converting logical descriptors to MEM 112 physical descriptors.

LD Bus 1902 is also connected to write input of Name Cache (NC) 1926. NC 1926 is a cache mechanism for storing logical descriptors corresponding to operand Names currently being used in user's programs. As previously described, Name Table Entries corresponding to operands currently being used in user's programs are stored in MEM 112. Certain Name Table Entries for operands of a user's program currently being executed are transferred from those Name Tables in MEM 112 to FU 120 and are therein evaluated to generate corresponding logical descriptors. These logical descriptors are then stored in NC 1926. As will be described further below, the instruction stream of a user's program is provided to FU 120's Instruction Buffer (IB) 1962 through MOD Bus 140. FU 120's Parser 1964 separates out, or parses, Names from IB 1962 and provides those Names as address inputs to NC 1924. NC 1924 in turn provides logical descriptor outputs to LD Bus 1902, and thus to input of ATU 1928. NC 1926 input from LD Bus 1902 allows logical descriptors resulting from evaluation of Name Table Entries to be written into NC 1926. FU 120's Protections Cache (PC) 1934 is a cache mechanism having an input connected from LD Bus 1902 and providing information, as described further below, regarding protection aspects of atear *i a cc a

C

i i i lc--- -82references to data in MEM 112 by user's programs. NC 1926, ATU 1928, and PC 1934 are thereby acceleration mechanisms of, respectively, CS 101's Namespace addressing, logical to physical address structure, and I 5 protection mechanism.

Referring again to DESP 1910, DESP 1910 includes BIAS 1952, connected from output of LENGRF 1936. As previously described, operands containing more than 32 data bits are transferred beteen MEM 112 and JP 114 by means of string transfers. In order to perform string transfers, it is necessary for FU 120 to generate a corresponding succession of logical descriptors wherein I .n length fields of those logical descriptors is no greater Sthan 5 bits, that is, specify lengths of no greater than S 15 32 data bits.

A logical descriptor describing a data item to be I transferred by means of a string transfer will be stored ,in GRF 506. AON field of the logical descriptor will reside in AONGRF 1932, O field in OFFGRF 1934, and L field in LENGRF 1936. At each successive transfer of a 32 bit word in the string transfer, O field of that original logical descriptor will be incremented by the number of data bits transferred while L field will be accordingly decremented. The logical descriptor residing in GRF 506 will thereby describe, upon each successsive i ttransfer of the string transfer, that portion of the data j t item yet to be transferred. 0 field in OFFGRF 1934 will indicate increasingly larger offsets into that data item, while L field will indicate successively shorter I w w -I I I -83lengths. AON and O fields of the logical descriptor in GRF 506 may be utilized directly as AON and 0 fields of the successive logical descriptors of the string transfer. L field of the logical descriptor residing in LENGRF 1936, however, may not be so used as L fields of the successive string transfer logical descriptors as this L field indicates remaining length of data item yet to be transferred. Instead, BIAS 1952 generates the bit L fields of successive string transfer logical descriptors while correspondingly decrementing L field of the logical descriptor in LENGRF 1936. During each transfer, BIAS 1952 generates L field of the neXt string transfer logical descriptor while concurrently providing L field of the current string transfer logical 15 descriptor. By doing so, BIAS 1952 thereby increases i speed of execution of string transfers by performing pipelined L field operations. BIAS 1952 thereby allows CS 101 to appear to the user to be a variable word length machine by automatically performing string transfers.

This mechanism is used for transfer of any data item greater than 32 bits, for example double precision floating point numbers.

Finally, FU 120 includes microcode circuitry for controlling all FU 120 operations described above. In particular, FU 120 includes a microinstruction sequence c control store (mC) 1920 storing sequences of re microinstructions for controlling step by step execution rt of all FU 120 operations. In general, these FU 120 h '1 I

I

-84operations fall into two classes. A first class includes those microinstruction sequences directly concerned with executing the SOPs of user's programs. The second class includes microinstruction sequences concerned with CS 101's operating systems, including and certain automatic, internal FU 120 functions such as evaluation of Name Table Entries.

As previously described, CS 101 is a multiple S- Language machine. For example, mC 1920 may contain microinstruction sequences for executing user's SOPs in at least four different Dialects. mC 1920 is comprised of a writeable control store and sets of microinstruction sequences for various Dialects may be transferred into and out of mC 1920 as required for execution of various 15 user's programs. By storing sets of microinstruction S" sequences for more than one Dialect in mC 1920, it is possible for user's programs to be written in a mixture of user languages. For example, a particular user's program may be written primarily in FORTRAN but may call certain COBOL routines. These COBOL routines will be '*correspondingly translated into .COBOL dialect SOPs and executed by COBOL microinstruction sequences stored in mC 1920.

S i The instruction stream provided to FU 120 from MEM 112 has been previously described with reference to Fig.

S3. SOPs and Names of this instruction stream are 1 t transferred from MOD Bus 140 into IB 1962 as they are provided from MEM 112. IB 1962 includes two 32 bit (one i j word) registers. IB 1962 also includes prefetch circuitry for reading for SOPs and Names of the instruction stream from MEM 112 in such a manner that IB 1962 shall always contain at least one SOPs or Name. FU 120 includes 1964 which reads and separates, or parses, SOPs and Names from IB 1962. As previously described, P 1964 provides those Names to NC 1926, which accordingly provides logical descriptors to ATU 1928 so as to read the corresponding operands from MEM 112.

SOPs parsed by P 1964 are provided as inputs to Fetch Unit Dispatch Table (FUDT) 1904 and Execute Unit Dispatch Table (EUDT) 1966. Referring first to FUDT 1904, FUDT 1904 is effectively a table for translating SOPs to j *starting addresses in mC 1912 of corresponding j 15 microinstruction sequences. This intermediate "f translation of SOPs to mC 1912 addresses allows efficient packing of microinstruction sequences within mC 1912.

That is, certain microinstruction sequences may be common to two or more S-Language Dialects. Such microinstruction sequences may therefore be written into S* mC 1912 once and may be referred to by different SOPs of different S-Language Dialects.

EUDT 1966 performs a similar function with respect to i EU 122. As will be described below, EU 122 contains a mC, similar to mC 1912, which is addressed through EUDT 1966 by SOPs specifying EU 122 operations. In addition, c^ FU 120 may provide such addresses mC 1912 to initiate EU r 122 operations as required to assist certain FU 120 -86operations. Examples of such operations which may be requested by FU 120 include calculations required in evaluating Name Table Entries to provide logical descriptors to be loaded into NC 1926.

Associated with both FUDT 1904 and EUDT 1966 are Dialect registers 1905 and 1967. D registers 1905 and 1967 store information indicating the particular S- Language Dialect currently being utilized in execution of a user's program. Outputs of D registers 1905 and 1967 are utilized as part of the address inputs to mC 1912 and EU 122's mC.

4. Execute--Unit-EU)+ 122 Fia.- 20)4 S As previously described, EU 122 is an arithmetic and logic unit provided to relieve FU 120 of certain arithmetic operations. EU 122 is capable of performing addition, subtraction, multiplication, and division operations on integer, packed and unpacked decimal, and single and double precision floating operands. EU 122 is an independently operating microcode controlled machine including Microcode Control (mC) 2010 which, as described 3 above, is addressed by EUDT 1966 to initiate EU 122 operations. mC 2010 also includes logic for handling mutual interrupts between FU 120 and EU 122. That is, FU t 120 may interrupt current EU 122 operations to call upon EU 122 to assist an FU 120 operation. For example, FU 120 may interrupt an arithmetic operation currently being i executed by EU 122 to call upon EU 122 to assist in generating a logical descriptor from a Name Table Entry.

-87- Similarly, EU 122 may interrupt current FU 120 operations when EU 122 requires FU 120 assistance in executing a current arithmetic operation. For example, EU 122 may interrupt a current FU 120 operation if EU 122 receives an instruction and operands requiring EU 122 to perform a divide by zero.

Referring to Fig. 20, a partial block diagram of EU 122 is shown. EU 122 includes two arithmetic and logic units. A first arithmetic and logic unit (MULT) 2014 is utilized to perform addition, subtraction, multiplication, and division operations upon integer and decimal operands, and upon mantissa fields of single and double precision floating point operands. Second ALU (EXP) 2016 is utilized to perform operations upon single 15 and double precision floating point operand exponent I' fields in parallel with operations performed upon floating point mantissa fields by MULT 2014. Both MULT S. 2014 and EXP 2016 include an arithmetic and logic unit, respectively MALU 2074 and EXPALU 2084. MULT 2014 and EXP 2016 also include register files, respectively MRF .1 2050 and ERF 2080, which operate and are addressed in parallel in a manner similar to AONGRF 1932, OFFGRF 1984 and LENGRF 1936.

2 5 E Operands for EU 122 to operate upon are provided from 25 MEM 112 through MOD Bus 140 and are transferred into Operand Buffer (OPB) 2022. In addition to serving as an i input buffer, OPB 2022 performs certain data format r I 1 1 -88manipulation operations to transform input operands into formats most efficiently operated with by EU 122. In particular, EU 122 and MULT 2014 may be designed to operate efficiently with packed decimal operands. OPB 2022 may transform unpacked decimal operands into packed decimal operands. Unpacked decimal operands are in the form of ASCII characters wherein four bits of each characters are binary codes specifying a decimal value between zero and nine. Other bits of each character are referred to as zone fields and in general contain information identifying particular ASCII characters. For example, zone field bits may specify whether a particular ASCII character is a number, a letter, or punctuation.

Packed decimal operands are comprised of a series of four l 15 bit fields wherein each field contains a binary number o 4 specifying a decimal value of between zero and nine. OPB 2022 converts unpacked decimal to packed decimal operands by extracting zone field bits and packing the four numeric value bits of each character into the four bit fields of a packed decimal number.

EU 122 is also capable of transforming the results of *i .arithmetic operands, for example in packed decimal format, into unpacked decimal format for transfer back to MEM 112 or FU 120. In this case, a packed decimal result appearing at output of MALU 2074 is written into MRF 2050 through a multiplexer, not shown in Fig. 20, which t transforms the four bit numeric code fields of the packed C C f -89decimal results into corresponding bits of unpacked decimal operand characters, and forces blanks into the zone field bits of those unpacked decimal characters.

The results of this operation are then read from MRF 2050 to MALU 2074 and zone field bits for those unpacked decimal characters are read from Constant Store (CST) 2060 to MALU 2074. These inputs from MRF 2050 and CST 2060 are added by MALU 2074 to generate final result outputs in unpacked decimal format. These final results may then be transferred onto JPD Bus 142 through Output Multiplexer (OM) 2024.

Considering first floating point operations, in 4*1. addition or subtraction of floating point operands it is It necessary to equalize the values of the floating point 15 operand exponent fields. This is referred to as prealignment. In floating point operations, exponent S fields of the two operands are transferred into EXPALU 2034 and compared to determine the difference between exponent fields. An output representing difference between exponent fields is provided from EXPALU 2034 to 4 an input of floating point control (FPC) 2002. FPC 2002 4* in turn provides control outputs to MALU 2074, which has received the mantissa fields of the two operands. MALU 2074, operating under direction of FPC 2002, accordingly right or left shifts one operand's mantissa field to effectively align that operand's exponent field with the other operand's exponent field. Addition or subtraction of the operand's mantissa fields may then proceed.

EXPALU 2034 also performs addition or subtraction of floating point operand exponent fields in multiplication or division operations, while MALU 2074 performs multiplication and division of the operand mantissa fields. Multiplication and division of floating point operand mantissa fields by MALU 2074 is performed by successive shifting of one operand, corresponding generation of partial products of the other operand, and successive addition and subtraction of those partial products.

Finally, EU 122 performs normalization of the results of floating point operand operations by left shifting of a final result's mantissa field to eliminate zeros in the most significant characters of the final result mantissa field, and corresponding shifting of the final result exponent fields. Normalization of floating point t ,operation results is controlled by FPC 2002. FPC 2002 examines an unnormalized floating point result output, of MALU 2074 to detect which, if any, of the most significant characters of that results contain zeros.

FPC 2002 then accordingly provides control outputs to 9..9.EXPALU 2034 and MALU 2074 to correspondingly shift the exponent and mantissa fields of those results so as to eliminate leading character zeros from the mantissa field. Normalized mantissa and exponent fields of floating point results may then be transferred from MALU 2074 and EXPALU 2034 to JPD Bus 142 through OM 2024.

rf c

A'}

-91- As described above, EU 122 also performs addition, subtraction, multiplication, and division operations on operands. In this respect, EU 122 uses a leading zero detector in FPC 2002 in efficiently performing multiplication and division operations. FPC 2002's leading zero detector examines the characters or bits of two operands to be multiplied or divided, starting from the highest, to determine which, if any, contain zeros so as not to require a multiplication or division' operation. FPC 2002 accordingly left shifts the operands to effectively eliminate those characters or bits, thus reducing the number of operatiohs to multiply or divide the operands and accordingly reducing the time required to operate upon the operands.

15 Finally, EU 122 utilizes a unique method, with associated hardware, foi performing arithmetic operations on decimal operands by utilizing circuitry which is otherwise conventionally used only to perform operations upon floating point operands. As described above, MULT 2074 is designed to operate with packed decimal operands, that is operands in the form of consecutive blocks of S- four bits wherein each block of four bits contains a binary code representing numeric values of between zero S and nine. Floating point operands are similarly in the form of consecutive blocks of four bits. Each block of no• four bits in a floating point operand, however, contains 4 a binary number representing a hexadecimal value of S1 i l -92between zero and fifteen. As an initial step in operating with packed decimal operandsp those operands are loaded, one at a time, into MALU 2074 and, with each such operand, a number comprised of all hexadecimal sixes is loaded into MALU 2074 from CST 2060. This CST 2060 number is added to each packed decimal operand to effectively convert those packed decimal operands into hexadecimal operands wherein the four bit blocks contain numeric values in the range of six to fifteen, rather than in the original range of zero to nine. MULT 2014 then performs arithmetic operation upon those transformed operands, and in doing so detects and saves information regarding which four bit characters of those operands have resulted in generation of carries during the arithmetic operations. In a final step, the intermediate m result resulting from completion of those arithmetic operations upon those transformed operands are reconverted to packed decimal format by subtraction of hexadecimal sixes from those characters for which carries t2 20 have been generated. Effectively, EU 122 converts packed I f, decimal operands into "Excess Six" operands, performs arithmetic operations upon those "Excess Six" operands, i and reconverts "Excess Six" results of those operations back into packed decimal format.

Finally, as previously descibed FU 120 controls c" transfer of arithmetic results from EU 122 to MEM 112.

In doing so, FU 120 generates a logical descriptor describing the size of MEM 112 address space, or L -93- "container", that result is to be transferred into. In certain arithmetic operations, for example integer operations, an arithmetic result may be larger than anticipated and may contain more bits than the MEM 112 "container". Container Size Check Circuit (CSC) 2052 compares actual size of arithmetic results and L fields of MEM 112 "container" logical descriptors, CSC 2052 generates an output indicating whether an MEM 112 "container" is smaller than an arithmetic result.

Having briefly described certain features of CS 101 structure and operation in the above overview, these and other features of CS 101 will be described in further detail next below in a more detailed introduction of CS 101 structure and operation. Then, in further 15 descriptions, these and other features of CS 101 t. structure and operation will be described in depth.

4 i. Introduction- Piags *lOl-.Q) A. General Structure and Operation-Fia.t iOV) Sa. General Structure 20 Referring to Fig. 101, a partial block diagram of Computer System (CS) 10110 is shown. Major elements of S• CS 10110 are Dual Port Memory (MEM) 10112, Job Processor (JP) 10114, Input/Output System (IOS) 10116, and Diagnostic Processor (DP) 10118. JP 10114 includes Fetch Unit (FU) 10120 and Execute Unit (EU) 10122.

e L 1 1 i -94- Referring first to IOS 10116, IOS 10116 is interconnected with External Devices (ED) 10124 through Input/Output Bus 10126. ED 10124 may include, for example, other computer systems, keyboard/display units, and disc drive memories. IOS 10116 is interconnected with Memory Input/Output (MIO) Port 10128 of MEM 10112 through Input/Output to Memory (IOM) Bus 10130 and Memory to Input/Output (MIO) Bus 10129, and with FU 10120 through I/O Job Processor (IOJP) Bus 10132.

DP 10118 is interconnected with, for example, external keyboard/CRT Display Unit (DU) 10134 through Diagnostic Processor Input/Output (DPIO) Bus 10136. DP 10118 is interconnected with IOS 10116, MEM 10112, FU 10120, and EU 10122 through Diagnostic Processor (DP) Bus 15 10138.

Ss" Memory to Job Processor (MJP) Port 10140 of Memory 10112 is interconnected with FU 10120 and EU 10122 through Job Processor Data (JPD) Bus 10142. An output of MJP 10140 is connected to inputs of FU 10120 and EU 10122 through Memory Output Data (MOD) Bus 10144. An output of FU 10120 is connected to an input of MJP 10140 through 9. Physical Descriptor (PD) Bus 10146. FU 10120 and EU 10122 are interconnected through Fetch/Execute Bus 10148.

b. Gener.l Operation As will be discussed further below, IOS 10116 and MEM 10112 operate independently under general control of JP 10114 in executing multiple user's programs. In this -bf t l 1 t regard, MEM 10112 is an intelligent, prioritizing memory having separate and independent ports MIO 10128 and MJP 10140 to IOS 10116 and JP 10114 respectively. MEM 10112 is the primary path for information transfer between External Devices 10124 (through IOS 10116) and JP 10114.

MEM 10112 thus operates both as a buffer for receiving and storing various individual user's programs data, instructions, and results of program execution) and as a main memory for JP 10114.

A primary function of IOS 10116 is as an input/output buffer between CS 10110 and ED 10124. Data and instructions are transferred from ED 10124 to IOS 10116 through I/O Bus 10126 in a manner and format compatible Sith ED 10124. IOS 10116 receives and stores this information, and manipulates the information into formats Ssuitable for transfer into MEM 10112. IOS 10116 then indicates to MEM 10112 that new information is available for transfer into MEM 10112. Upon acknowledgement by MEM 10112, this information is transferred into MEM 10112 through IOM Bus 10130 and MIO Port 10128. MEM 10112 stores the information in selected portions of MEM 10112 tphysical address space. At this time, IOS 10116 notifies JP 10114 that new information is present in MEM 10112 by providing a "semaphore" signal to FU 10120 through IOJP Bus 10132. As will be described further below, CS 10110 manipulates the data and instructions stored in MEM 10112 into certain information structures used in executing user's programs. Among these structures are certain 4- i l -96structures, discussed further below, which are used by CS 10110 in organizing and controlling flow and execution of user programs.

FU 10120 and EU 10122 are independently operating microcode controlled "machines" together comprising the CS 10110 micromachine for executing user's programs stored in MEM 10112. Among the principal functions of FU 10120 are: fetching and interpreting instructions and data from MEM 10112 for use by FU 10120 and EU 10122; organizing and controlling flow of user programs; initiating EU 10122 operations; (4) performing arithmetic and logic operations on data; rt controlling transfer of data from FU 10120 and EU 10122 r €to MEM 10112; and, maintaining certain "stack" and "register" mechanisms, described below. FU 10120 "cache" It t t t k 15 mechanisms, also described below, are provided to enhance the speed of operation of JP 10114. These cache mechanisms are acceleration circuitry including, in part, high speed memories for storing copies of selected information stored in MEM 10112. The information stored t ZO20 in this acceleration circuitry is therefore more rapidly available to JP 10114. EU 10122 is an arithmetic unit capable of executing integer, decimal, or floating point arithmetic operations. The primary function of EU 10122 is to relieve FU 10120 from certain extensive arithmetic operations, thus enhancing the efficiency of CS 10110.

\j I g. .1CI L. 3 1 97- In general, operations in JP 10114 are executed on a memory to memory basis; data is read from MEM 10112, operated upon, and the results returned to MEM 10112. In this regard, certain stack and cache mechanisms in JP 10114 (described below) operate as extensions of MEM 10112 address space.

In operation, FU 10120 reads data and instructions from MEM 10112 by providing physical addresses to MEM 10112 by way of PA Bus 10146 and MJP Port 10140. The instructions and data are transferred to FU 10120 and EU 10122 by way of MJP Port 10140 and MOD Bus 10144.

Instructions are interpreted by FU 10120 microcode circuitry, not shown in Fig. 101 but described below, and when necessary, microcode instructions are provided to EU S 15 10122 from FU 10120's microcode control by way of F/E Bus 10148, or by way of JPD Bus 10142.

qu' As stated above, FU 10120 and EU 10122 operate f asynchronously with respect to each other's functions. A microinstruction from FU 10120 microcode circuitry to EU 10122 may initiate a selected operation of EU 10122. EU t u, 10122 may then proceed to independently execute the selected operation. FU 10120 may proceed to concurrently execute other operations while EU 10122 is completing the St, selected arithmetic operation. At completion of the selected arithmetic operation, EU 10122 signals FU 10120 that the operation results are available by way of a "handshake" signal through F/E Bus 10148. FU 10120 may 4' y^WW^ -98then receive the arithmetic operation results for further processing or, as discussed momentarily, may directly transfer the arithmetic operation results to MEM 10112.

As described further below, an instruction buffer referred to as a "queue" between FU 10120 and EU 10122 allows FU 10120 to assign a sequence of arithmetic operations to be performed by EU 10122.

Information, such as results of executing an instruction, is written into MEM 10112 from FU 10120 or EU 10122 by way of JPD Bus 10142. FU 10120 provides a "physical write address" signal to MEM 10112 by way of PA Bus 10146 and MJP Port 10140. Concurrently, the information to be written into MEM 10112 is placed on JPD Bus 10142 and is subsequently written into MEM 10112 at the locations selected by the physical write address.

FU 10120 places a semaphore signal on IOJP Bus 10132 Sto signal to IOS 10116 that information, such as the results of executing a user's program, is available to be read out of CS 10110. IOS 10116 may then transfer the 20 information from MEM 10112 to IOS 10116 by way of MIO Port 10128 and IOM Bus 10130. Information stored in IOS 10116'is then transferred to ED 10124 through I/O Bus 10126.

During execution of a user's program, certain information required by JP 10116 may not be available in Sa MEM 10112. In such cases as further described in a following discussion, JP 10114 may write a request for information into MEM 10112 and notify IOS 10116, by way I i -99- 4**4 4 4e '.94 ft Ii f f 4r f 9 9# 4 444 4 9.

4 9 4449.J 44 94 9i $4 4 4.' of IOJP Bus 10132, that such a request has been made.

IOS 10116 will then read the request and transfer the desired information from ED 10124 into MEM 10112 through IOS 10116 in the manner described above. In such operations, IOS 10116 and JP 10114 operate together as a memory manager wherein the memory space addressable by JP 10114 is termed virtual memory space, and includes both MEM 10112 memory space and all external devices to which IOS 10116 has access.

As previously described, DP 10118 provides a second interface between Computer System 10110 and the external world by way of DPIO Bus 10136. DP 10118 allows DU 10134, for example a CRT and keyboard unit or a teletype, to perform all functions which are conventionally provided by a hard switches and lights) console.

For example, DP 10118 allows DU 10134 to exercise control of Computer System 10110 for such purposes as system initialization and start up, execution of diagnostic processes, and fault monitoring and identification. DP 20 10118 has read and write access to most memory and register portions within each of IOS 10116, MEM 10112, FU 10120, and EU 10122 by way of DP Bus 10138. Memories and registers in CS 10110 can therefore be directly loaded or initialized during system start up, and can be directly 25 read or loaded with test and diagnostic signals for fault monitoring and identification. In addition, as described i hi r j r

~T

aii -100further below, microinstructions may be loaded into JP 10114's microcode circuitry at system start up or as required.

Having described the general structure and operation of Computer System 10110, certain features of Computer System 10110 will next be briefly described to aid in understanding the following, more detailed descriptions of these and other features of Computer System 10110.

c. Definition-, of.Certa-n--Terms Certain terms are used relating to the structure and operation of CS 10110 throughout the following discussions. Certain of these terms will be discussed and defined first, to aid in understanding the following descriptions. Other terms will be introduced in the following descriptions as required.

A ro edure. is a sequence of operational steps, or instructions, to be executed to perform some operation.

A procedure may include data to be operated upon in performing the operation.

A oratam is a static group'of one or more procedures. In general, programs may be classified as user programs, utility programs, and operating system programs. A user program is a group of procedures generated by and private to one particular user of a group of users interfacing with CS 10110. Utility r. r ~programs are commonly available to all users; for example, a compiler comprises of a set of procedures for compiling a user language program into an S-language ria~p-:-IL.1 s 8 1 -101-

SI;

S. 5$ 9

S

4.r r* 4* 4 I S S 4r 5 program. Operating system programs are groups of procedures internal to CS 10110 for allocation and control of CS 10110 resources. Operating system programs also define interfaces within CS 10110. For example, as will be discussed further below all operands in a program are referred to by "NAME". An operating system program translates operand NAME into the physical locations of the operands in MEM 10112. The NAME translation program thus defines the interface between operand NAME (name space addresses) and MEM 10112 physical addresses.

A process is an independent locus of control passing through physical, logical or virtual address spaces, or, more particularly, a path of execution through a series of programs procedures). A process will generally include a user program and data plus one or more utility programs a compiler) and operating system programs necessary to execute the user program.

An bject is a uniquely identifiable portion of "data space" accessible to CS 10110. An object may be regarded as a container for information and may contain data or procedure information or both. An object may contain for example, an entire program, or set of procedures, or a single bit of data. Objects need not be contiguously located in the data space accessible to CS 10110, and the information contained in an object need not be contiguously located in that object.

i -102- A domain is a state of operation of CS 10110 for the purposes of CS 10110's protection mechanisms. Each domain is defined by a set of procedures having access to objects within that domain for their execution. Each object has a single domain of execution in which it is executed if it is a procedure object, or used, if it is a data object. CS 10110 is said to be operating in a particular domain if it is executing a procedure having that domain of execution. Each object may belong to one or more domains; an object belongs to a domain if a procedure executing in that domain has potential access to the object. CS 10110 may, for example have four domains: User domain, Data Base Management System (DBMS) t domain, Extended Operating System (EOS) domain, and Kernel Operating System (KOS) domain. User domain is the domain of execution of all user provided procedures, such as user or utility procedures. DBMS domain is the domain 4 *4 of execution for operating system procedures for storing, retrieving, and handling data. EOS domain is the domain 20 of execution of operating system procedures defining and forming the user level interface with CS 10110, such as procedures for controlling an executing files, processes, and I/O operations. KOS domain is the domain of execution of the low level, secure operating system which manages and controls CS 10110's physical resources.

Other embodiments of C6 10110 may have fewer or more domains than those just described. For example, DBMS procedures may be incorporated into the EOS domain or EOS 4. 1 -103domain may be divided by incorporating the I/O procedures into an I/O domain. There is no hardware enforced limitation on the number of, of boundaries between, domains in CS 10110. Certain CS 10110 hardware functions and structures are, however, dependent upon domains.

A subject is defined, for purposes of CS 10110's protection mechanisms, as a combination of the current principle (user), the current process being executed, and the domain the process is currently being executed in.

In addition to principle, process, and domain, which are identified by UIDs, subject may include a Tag, which is a user assigned identification code used where added t 4t security is required. For a given process, principle and process are constant but the domain is determined by the 4i 15 procedure currently being executed. A process's associated subject is therefore variable along the path of execution of the process.

Having discussed and defined the above terms, certain features of CS 10110 will next be briefly described.

220 d. Multi-Prog-ram--Oeration CS 10110 is capable of concurrently executing two or more programs and selecting the sequence of execution of S*1' programs to make most effective use of CS 10110's resources. This is referred to as multiprogramming. In this regard, CS 10110 may temporarily suspend execution Sof one program, for example when a resource or certain information required for that program is not immediately available, and proceed to execute another program until N i W W iilii^ y i 1 1 -104the required resource or information becomes available.

For example, particular information required by a first program may not be available in MEM 10112 when called for. JP 10114 may, as discussed further below, suspend execution of the first program, transfer a request for that information to IOS 10116, and proceed to call and execute a second program. IOS 10116 would fetch the requested information from ED 10124 and transfer it into MEM 10112. At some time after IOS 10116 notifies JP 10114 that the requested information is available in MEM 10112, JP 10114 could suspend execution of the second program and resume execution of the first program.

e. Multi-Lanuage--OpenraTion N 44 As previously described, CS 10110 is a multiple 15 language machine. Each program written in a high level .e user language, such as COBOL or FORTRAN, is compiled into Sa corresponding Soft Language program. That is, in terms of a conventional computer system, each user level language has a corresponding machine language, il, 20 classically defined as an assembly language. In contrast t\ to classical assembly languages, S-Languages are mid- S" level languages wherein each command in a user's high 4tI. level language is replaced by, in general, two or three S-Language instructions, referred to as SINs. Certain SINs may be shared by two or more high level user languages. CS 101,10, as further described in following discussions, provides a set, or dialect, of microcode ir instructions (S-Interpreters) for each S-Language. S- Interpreters interpret SINs and provide corresponding r -105sequences of microinstructions for detailed control of CS 10110. CS 10110's instruction set and operation may therefore be tailored to each user's program, regardless of the particular user language, so as to most efficiently execute the user's program. Computer System 10110 may, for example, execute programs in both FORTRAN and COBOL with comparable efficiency. In addition, a user may write a program in more than one high level user language without loss of efficiency. For example, a user may write a portion of his program in COBOL, but may wish to write certain portions in FORTRAN. In such cases, the COBOL portions would be compiled into COBOL SINs and executed with the COBOL dialect S-Interpreter. The t L FORTRAN portions would be compiled into FORTRAN SINs and It t1f 41 executed with a FORTRAN dialect S-Interpreter. The 15 present embodiment of CS 10110 utilizes a uniform format for all SINs. This feature allows simpler S-Interpreter structures and increases efficiency of SIN interpretation because it is not necessary to provide means for interpreting each dialect individually.

20 f. Addressing--Stfucture Each object created for use in, or by operation of, a S. CS 10110 is permanently assigned a Unique Identifier (UID). An object's UID allows that object to be uniquely Sidentified and located at any time, regardless of which particular CS 10110 it was created by or for or where it is subsequently located. Thus each time a new object is defined, a new and unique UID is allocated, much as

I

pr -106t t #4 i *I t

I

1*44 F* a *q 4 44 9.

*e I C I c' e social security numbers are allocated to individuals. A particular piece of information contained in an object may be located by a logical address comprising the object's UID, an offset from the start of the object of the first bit of the segment, and the length (number of bits) of the information segment. Data within an object may therefore be addressed on a bit granular basis. As will be described further in following discussions, UID's are used within a CS 10110 as logical addresses, and, for example, as pointers. Logically, all addresses and pointers in CS 10110 are UID addresses and pointers. As previously described and as described below, however, short, temporary unique identifiers, valid only within JP 10114 and referred to as Active Object Numbers are used 15 within JP 10114 to reduce the width of address buses and amount of address information handled.

An object becomes active in CS 10110 when it is transferred from backing store CED 10124 to MEM 10112 for use in executing a process. At this time, each such 20 object is assigned an Active Object Number (AON). AONs are short unique identifiers and are related to the object's UIDs through certain CS 10110 information structures described below. AONs are used only within JP 10114 and are used in JP 10114, in place of UIDs, to reduce the required width of JP 10114's address buses and the amount of address data handled in JP 10114. As with UID logical addresses, a piece of data in an object may be addressed through a bit granular AON logical address comprising the object's AON, an offset from the start of dA AM4 -107the object of the first bit of the piece, and the length of the piece.

The transfer of logical addresses, for example pointers, between MEM 10112 (UIDA) and JP 10114 (AONs) during execution of a process requires translations between UIDs and AONs. As will be described in a later discussion, this translation is accomplished, in part, through the information structures mentioned above.

Similarly, translation of logical addresses to physical addresses in MEM 10112, to physically access information stored in MEM 10112, is accomplished through CS 10110 information structures relating AON logical addresses to r* l t MEM 10112 physical addresses.

Each operand appearing in a program is assigned a 15 Name when the program is compiled. Thereafter, all references to the operands are through their assigned Names. As will be described in detail in a later discussion, CS 10110's addressing structure includes a mechanism for recognizing Names as they appear in an S 20 instruction stream and Name Tables containing directions Sfor resolving Names to AON logical addresses. AON logical addresses may then be evaluated, for example translated into a MEM 10112 physical address, to provide actual operands. The use of Names to identify operands in the instructions stream (process) allows a Scomplicated address to be replaced by a simple reference of uniform format; does not require that an operation be directly defined by data type in the instruction streaml allows repeated references to an operand to be made in an instruction stream by merely L ;i -108repeating the operand's Name; and, allows partially completed Name to address translations to be stored in a cache to speed up operand references. The use of Names thereby substantially reduces the volume of information required in the instruction stream for operand references and increases CS 10110 speed and efficiency by performing operands references through a parallel operating, underlying mechanism.

Finally, CS 10110 address structure incorporates a i 0 set of Architectural Base Pointers (ABPs) for each process. ABPs provide an addressing framework to locate data and procedure information belonging to a process and are used, for example, in resolving Names to AON logical addresses.

g. -rotecti on- -Mechanism CS 10110's protection mechanism is constructed to prevent a user from gaining access to or disrupting another user's process, including data, and (2) interfering with or otherwise subverting the operation of Co 10110. Access rights to each particular active object are dynamically granted as a function of the currently active subject. A subject is defined by a combination of the current principle (user), the current process being executed, and the domain in which the process is currently being executed. In addition to principle, c !process, and domain, subject may include a Tag, which is a user assigned identification code used where added security is required. For a given process, the principle -109and process are constant but the domain is determined by the procedure currently being executed. A process's associated subject is therefore variable along the path of execution of the process.

In a present embodiment of CS 10110, procedures having KOS domain of execution have access to objects in KOS, EOS, DBMS, and User domains; procedures having EOS domain of execution have access to objects in EOS, DBMS, and User domains; procedures having DBMS domain of execution have access to objects in DBMS and User domains; and procedures having User domain of execution have access only to objects in User domain. A user cannot, therefore, obtain access to objects in KOS domain of execution and cannot influence CS 10110's low level, Of 15 secure operating system. The user's process may, S* however, call for execution a procedure having KOS domain of execution. At this point the process's subject is in the KOS domain and the procedure will have access to certain objects in KOS domain.

In a present embodiment of CS 10110, also described in a later discussion, each object has associated with it an Access Control List (ACL) An ACL contains an Access Control Entry (ACE) for each subject having access to •""that object. ACEs specify, for each subject, access rights a subject has with regard to that object.

There is normally no relationship, other than that l defined by an object's ACL, between subjects and objects. CS 10110, however, supports Extended Type -110- Objects having Extended ACLs wherein a user may specifically define which subjects have what access rights to the object.

In another embodiment of CS 10110, described in a following discussion, access rights are granted on a dynamic basis. In executing a process, a procedure may call a second procedure and pass an argument to the called procedure. The calling procedure will also pass selected access rights to that argument to the called procedure. The passed access rights exist only for the duration of the call.

In the dynamic access embodiment, access rights are granted only at the time they are required. In the ACL t «embodiment, access rights are granted upon object creation or upon specific request. In either embodiment, each procedure to which arguments may be passed in a cross-domain call has associated with it an Access r:J* "Information Array (AIA). A procedure's AIA states what access rights a calling procedure (subject) must have before the called procedure can operate on the passed argument. CS 10110's protection mechanisms compare the calling procedure's access rights to the rights required by the called procedure. This ensures that a calling Sprocedure may not ask a called procedure to do what the calling procedure is not allowed to do.' Effectively, a calling procedure can pass to a called procedure only the i access rights held by the calling procedure.

I

-111- Having described the general structure and operation and certain features of CS 10110, those and other features of CS 10110 operation will next be described in greater detail.

B. Computer -System--10-11l0--nformation-Structtes--and Mechanisms--Fgqs-. CS 10110 contains certain information structures and mechanisms to assist in efficient execution of processes. These structures and mechanisms may be considered as falling into three general types. The first type concerns the processes themselves, i.e., procedure and data objects comprising a user's process or directly related to execution of a user's process. The t r second type are for management, control, and execution of 15 processes. These structures are generally shared by all processes active in CS 10110. The third type are CS 10110 micromachine information structures and mechanisms. These structures are concerned with the eternal operation of the CS 10110 micromachine and are private to the CS 10110 micro-machine.

-Introduction-- F -ig 10-2 Referring to Fig. 102, a pictorial representation of CS 10110 (MEM 10112, FU 10120, and EU 10122) is shown i with certain information structures and mechanisms depicted therein. It should be understood that these information structures and mechanisms transcend or "cut across" the boundaries between MEM 10112, FU 10120, EU 10122, and IOS 10116. Referring to the upper portion of Fig. 103 Process Structures 10210 contains those -112information structures and mechanisms most closely concerned with individual processes, the first and third types of information structures described above. Process Structures 10210 reside in MEM 10112 and Virtual Processes 10212 include Virtual Processes (VP) 1 through N. Virtual Processes 10212 may contain, in a present embodiment of CS 10110, up to 256 VP's. As previously described, each VP includes certain objects particular to a single user's process, for example stack objects previously described and further described in a following description. Each VP also includes a Process Object containing certain information required to execute the Sprocess, for example pointers to other process information.

Virtual Processor State Blocks (VPSBs) 10218 include e" VPSBs containing certain tables and mechanisms for managing execution of VPs selected for execution by CS S• 10110.

A particular VP is bound into CS 10110 when a Virtual Process Dispatcher, described in a following discussion selects that VP as eligible for execution. The selected VPs Process Object, as previously described, is swapped into a VPSB. VPSBs 10218 may contain, for example 16 or 32 State Blocks so that CS 10110 may concurrently execute up to 16 or 32 VPs. then a VP assigned to a VPSB is to be executed, the VP is swapped onto the information structures and mechanisms shown in FU 10120 and EU 10122. FU Register and Stack Mechanism (FURSM) 10214 and EU Register and Stack Mechanism (EURSM) 10216, shown -112information structures and mechanisms most closely concerned with individual processes, the first and third types of information structures described above. Process Structures 10210 reside in MEM 10112 and Virtual Processes 10212 include Virtual Processes (VP) 1 through N. Virtual Processes 10212 may contain, in a present embodiment of CS 10110, up to 256 VP's. As previously described, each VP includes certain objects particular to a single user's process, for example stack objects previously described and further described in a following description. Each VP also includes a Process Object containing certain information required to execute the j .process, for example pointers to other process ,information.

Virtual Processor State Blocks (VPSBs) 10218 include VPSBs containing certain tables and mechanisms for r managing execution of VPs selected for execution by CS 10110.

A particular VP is bound into CS 10110 when a Virtual Process Dispatcher, described in a following discussion |H s *selects that VP as eligible for execution. The selected VPs Process Object, as previously described, is swapped into a VPSB. VPSBs 10218 may contain, for example 16 or t 32 State Blocks so that CS 10110 may concurrently execute up to 16 or 32 VPs. When a VP assigned to a VPSB is to be executed, the VP is swapped onto the information structures and mechanisms shown in FU 10120 and EU 10122. FU Register and Stack Mechanism (FURSM) 10214 and EU Register and Stack Mechanism (EURSM) 10216, shown -113respectively in FU 10120 and EU 10122, comprise register and stack mechanisms used in execution of VPs bound to CS 10110. These register and stack mechanisms, as will be discussed below, are also used for certain CS 10110 process management functions. Procedure Objects (POs) 10213 contains Procedure Objects (POs) 1 to N of the processes executing in CS 10110.

Addressing Mechanisms (AM) 10220 are a part of CS 10110's process management system and are generally associated with Computer System 10110 addressing functions as described in following discussions. UID/AON Tables 10222 is a structure for relating UID's and AON's, previously discussed. Memory Management Tables 10224 includes structures for relating AON logical E 15 addresses and MEM 10112 physical addresses; managing t r.

MEM 10112's physical address space; managing .transfer of information between MEM 10112 and CS 10110's t backing store (ED 10124) and, activating objects into CS 10110; Name Cache (NC) 10226 and Address Translation Cache (ATC) 10228 are acceleration mechanisms for storing addressing information relating to the VP currently bound to CS 10110. NC 10226, described further below, contains information relating operand Names to AON addresses. ATC 10228, also discussed further below, contains information relating AON addresses to MEM 10112 physical addresses. t.

r J W. -114- Protection Mechanisms 10230, depicted below AM 10220, include Protection Tables 10232 and Protection Cache (PC) 10234. Protection Tables 10232 contain information regarding access rights to each object active in CS 10110. PC 10234 contains protection information relating to certain objects of the VP currently bound to CS 10110.

Microinstruction Mechanisms 10236, depicted below PM 10230, includes Micro-code (M Code) Store 10238, FU (Micro-code) M Code Structure 10240, and EU Micro-code (M Code) Structure 10242. These structures contain microinstruction mechanisms and tables for interpreting SINs and controlling the detailed operation of CS 10110.

.Micro-instruction Mechanisms 10232 also provide microcode tables and mechanisms used, in part, in operation of the low level, secure operating system that manages and controls CS 10110's physical resources.

Having thus briefly described certain CS 10110 information structures and mechanisms with the aid of Fig. 102, those information structures and mechanisms will next be described in further detail in the order mentioned above. In these descriptions it should be noted that, in representation of MEM 10112 shown in Fig.

102 and in other figures of following discussions, the addressable memory space of MEM 10112 is depicted. I Certain portions of MEM 10112 address space have been t designated as containing certain information structures and mechanisms. These structures and mechanisms have real physical existence in MEM 10112, but may vary in both location and volume of MEM 10112 address space they occupy. Assigning position of a single, large memory to 1 i i' r ii-~i~ j' -L CI~I~ I t It *~c I I*

II

*51*( 9 Ir #5 -115contain these structures and mechanisms allows these structures and mechanisms to be reconfigured as required for most efficient operation of CS 10110. In an alternate embodiment, physically separate memories may be used to contain the structures and mechanisms depicted in MEM 10112, rather than assigned portions of a single memory.

b. Process- Structues-1-O-210-- FPis-.----3-14- Referring to Fig. 103, a partial schematic representation of Process Structures 10210 is shown.

Specifically, Fig. 103 shows a Process 10310 selected for execution, and its associated Procedure Objects (POs) in Process Objects (POs) 10213. P 10310 is represented 15 in Fig. 103 as including four procedure objects in POs 10213. It is to be understood that this representation is for clarity of presentation; a particular P 10310 may include any number of procedure objects. Also for clarity of presentation, EURSM 10216 is not shown as EURSM 10216 is similar to FURSM 10214. EURSM 10216 will be described in detail in the following detailed discussons of CS 10110's structure and operation.

As previously discussed, each process includes certain data and procedure object. As represented in Fig. 103 for P 10310 the procedure objects reside in POs 10213. The data.objects include Static Data Areas and stack mechanisms in P 10310. POs, for example KOS Procedure Object (KOSPO) 10318, contain the various I JF i z tr -116procedures of the process, each procedure being a sequence of SINs defining an operation to be performed in executing the process. As will be described below, Procedure Objects also contain certain information used in executing the procedures contained therein. Static Data Areas (SDAs) are data objects generally reserved for storing data having an existence for the duration of the process. P 10310's stack mechanisms allow stacking of procedures for procedure calls and returns and for swapping processes in and out of JP 10114. Macro-Stacks (MAS) 10328 to 10334 are generally used to store automatic data (data generated during execution of a procedure and having an existence for the duration of that procedure). Although shown as separate from the stacks in P 10310, the SDAs may be contained with MASs 10328 to 10334. Secure Stack (SS) 10336 stores, in general, CS 10110 micro-machine state for each procedure Scalled. Information stored in SS 10336 allows machine state to be recovered upon return from a called procedure, or when binding (swapping) a VP into CS 10110.

As shown in P 10310, each process is structured on a domain basis. A P 10310 may therefore include, for each I domain, one or more procedure objects containing 3 procedures having that domain as their domain of execution, an SDA and an MAS. For example, KOS domain of P 10310 includes KOSPO 10318, KOSSDA 10326, and KOSMAS 10334. P 10310's SS 10336 does not reside in any single domain of P 10310, but instead is a stack mechanism belonging to CS 10110 micromachine.

44 i -117- Having described the overall structure of a P 10310, the individual information structures and mechanisms of a P 10310 will next be described in greater detail.

1. Procedufe--Objects--&Piqt.--l& KOSPO 10318 is typical of CS 10110 procedure objects and will be referred to for illustration in the following discussion. Major components of KOSPO 10318 are Header 10338, External Entry Descripter (EED) Area 10340, Internal Entry Descripter (IED) Area 10342, S-Op Code Area 10344, Procedure Environment Descripter (PED) 10348, Name Table (NT) 10350, and Access Information Array (AIA) Area 10352.

Header 10338 contains certain information identifying PO 10318 and indicating the number of entries in EED area 10340, discussed momentarily.

t EED area 10340 and IED area 10342 together contain an Entry Descripter (ED) for each procedure in KOSPO 10318.

fll. KOSPO 10318 is represented as containing Procedures 1, 2, and 11, of which Procedure 11 will be used as an example in the present discussion. EIs effectively comprise an index through certain all information in KOSPO 10318 can be located. IEDs form an index to all KOSPO 10318 procedures which may be called only from other procedures 47 contained in KOSPO 10318. EEDs form an index to all KOSPO 10318 procedures which may be called by procedures F. external to KOSPO 10318. Externally callable procedures

S

i -118are distinguished aid, as described in a following discussion of CS 10110's protection mechanisms, in confirming external calling procedure's access rights.

Referring to ED 11, ED for procedure 11, three fields are shown therein. Procedure Environment Descripter Offset (PEDO) field indicates the start, relative to start of KOSPO 10318, of Procedure 11's PED in PED Area 10348. As will be discussed further below, a procedure's PED contains a Set of pointers for locating information used in the execution of that procedure. PED Area 10348 contains a PED for each procedure contained in 10318. In the present embodiment of CS 10110, a single PED may be shared by two or more procedures. Code Entry Point (CEP) field indicates the start, relative to Procedure Base S 15 Pointer (PBP) which will be discussed below, of Procedure t t 11's SIN Code and SIN Code Area 10344. Finally, ED 11's Initial Frame Size (IFS) field indicates the required .Initial Frame Size of the KOSMAS 10334 frame storing Procedure 11's automatic data.

PED 11, Procedure 11's PED in PED Area 10348, contains a set of pointers for locating information used Sin execution of Procedure 11. The first entry in PED 11 is a header containing information identifying PED 11.

i PED 11's Procedure Base Pointer (PBP) entry is a pointer providing a fixed reference from which other information in PO 10318 may be located. In a specific example, Procedure 11's CEP indicates the location, relative to -119- PBP, of the start of Procedure 11's S-Op code in S-Op Code Area 10344. As will be described further below, PBP is a CS 10110 Architectural Base Pointer (ABP). CS 10110's ABP's are a set of architectural pointers used in CS 10110 to facilitate addressing of CS 10110's address space. PED 11's Static Data Pointer (SDP) entry points to data, in PO 10318, specifying certain parameters of P 10310's KOSSDA 10326. Name Table Pointer (NTP) entry is a pointer indicating the location, in NT 10350, of Name Table Entry's (NTE's) for Procedure 11's operands. NT 10350 and NTE's will be described in greater detail in the following discussion of Computer System 10110's Addressing Structure. PED 11's S-Interpreter Pointer (SIP) entry is a pointer, discussed in greater detail in a following discussion of CS 10110's microcode structure, pointing to the particular S-Interpeter (SINT) to be used in interpreting Procedure 11's SIN Code.

i Referring finally to AIA 10352, AIA 10352 contains, as previously discussed, information pertaining to access rights required of any external procedure calling a 10318 ;procedure. There is an AIA 10352 entry for each PO 10318 procedure which may be called by an external procedure.

A particular AIA entry may be shared by one or more procedures having an ED in EED Area 10340. Each EED contains certain information, not shown for clarity of presentation, indicating that that procedure's -120corresponding AIA entry must be referred to, and the calling procedure's access rights confirmed, whenever that procedure is called.

2. Stack" Mechan sms--4Pi+ -l94-.-l5 As previously described, P 10310's stack mechanisms include SS 10336, used in part for storing machine state, and MAS's 10328'to 10334, used to store local data generated during execution of P 10310's procedures. P 10310 is represented as containing an MAS for each CS 10110 domain. In an alternate embodiment of CS 10110, a particular P 10310 will include MAS's only for those domains in which that P 10310 is executing a procedure.

Referring to MAS's 10328 to 10334 and SS 10336, P 10310 is represented as having had eleven procedure calls. Procedure 0 has called Procedure 1, Procedure 1 Sr* has called Procedure 2, and so on. Each time a procedure .is called, a corresponding stack frame is constructed on the MAS of the domain in which the called procedure is executed. For example, Procedures 1, 2, and 11 execute in KOS domain; MAS frames for Procedures 1, 2, and 11 therefore are placed on KOSMAS 10334. Similarly, Procedures 3 and 9 execute in EOS domain, so that their stack frames are placed on EOSMAS 10332. Procedures and 6 execute in DBMS domain, so that their stack frames 25 are placed on DBMSMAS 10330. Procedures 4, 7, 8, and execute in User domain with their stack frames being placed on USERMAS 10328. Procedure 11 is the most -121recently called procedure and procedure 11's stack frame on KOSMAS 10334 is referred to as the current frame.

Procedure 11 is the procedure which is currently being executed when VP 10310 is bound to CS 10110.

SS 10336, which is a stack mechanism of CS 10110 micromachine, contains a frame for each of Procedures 1 to 11. Each SS 10336 frame contains, in part, CS 10110 operating state for its corresponding procedure.

Referring to Fig. 104, a schematic representation of a typical MAS, for example KOSMAS 10334, is shown.

KOSMAS 10334 includes Stack Header 10410 and a Frame 10412 for each procedure on KOSMAS 10334. Each Frame 10412 includes a Frame Header 10414, and may contain a Linkage Pointer Block 10416, a Local Pointer Block 10418, and a Local (Automatic) Data Block 10420.

KOSMAS 10334 Stack Header 10410 contains at least the following information: an offset, relative to Stack Header 10410, indicating the location of Frame Header 10414 of the first frame on KOSMAS 10334; a Stack Top Offset (STO) indicating location, relative to start of KOSMAS 10334, of the top of KOSMAS 10334; top of KOSMAS 10334 is indicated by pointer STO Spointing to the top of the last entry of Procedure 11 I +25 Frame 10412's Local Data Block 10420; an offset, relative to start of KOSMAS 10334, indicating location of Frame Header 10414 of the current top frame of KOSMAS 10334; in Fig. 104 this offst. s iL -122represented by Frame Pointer an ABP discussed further below; the VP 10310's UID; a UID pointer indicating location of certain domain environment information, described further in a following discussion; a signaller pointer indicating the location of certain routines for handling certain CS 10110 operating system faults; a UID pointer indicating location of KOSSDA 10326; and a frame label sequencer containing pointers to headers of frames in other domains; these pointers are used in executing non-local go-to operations.

S 15 KOSMAS 10334 Stack Header 10410 thereby contains B information for locating certain important points in KOSMAS 10334's structure, and for locating certain \information pertinent to executing procedures in KOS domain.

Each Frame Header 10414 contains at least the following information: S. offsets, relative to the Frame Header 10414, indicating the locations of Frame Headers 10414 of the S*previous and next frames of KOSMAS 10334; an offset, relative to the Frame Header 10414, v ~indicating the location of the top of that Frame 10412; information indicating the number of passed arguments contained in that Frame 10412; '1o -123a dynamic back pointer, in UID/Offset format, to the previous Frame 10412 if that previous Frame 10412 resides in another domain; a UID/Offset pointer to the environmental descripter of the procedure calling that procedure; a frame label sequence containing information indicating the locations of other Frame Headers 10414 in KOSMAS 10334; this information is used to locate other frames in KOSMAS 10334 for the purpose of executing local go-to operations. Frame Headers 10414 thereby contain information for locating certain important points in KOSMAS 10334 structure, and certain data pertinent to executing the associated procedures. In addition, Frame Headers 10414, in combination with Stack Header 10410, contain information for linking the activation records of each VP 10310 MAS, and for linking together the S)T: activation records of the individual MAS's.

4 Linkage Pointer Blocks 10416 contain pointers to arguments passed from a calling procedure to the called procedure. For example, Linkage Pointer Block 10416 of S.tt 2 Procedure 11's Frame 10412 will contain pointers to arguments passed to Procedure 11 from Procedure 10. The use of linkage pointers in CS 10110's addressing structure will be discussed further in a following discussion of CS 10110's Addressing Structure. Local Data Pointer Blocks 10418 contain pointers to certain of Si the associated procedure's local data. Indicated in Fig.

104 is a pointer, Frame Pointer pointing between l 4 1 i_ :Illp-- RUII ICCIIIC-- -124top most Frame 10412's Linkage Pointer Block 10416 and Local Data Pointer Block 10418. FP, described further in following discussions, is an ABP to MAS Frame 10412 of the process's current procedure.

Each Frame 10412's Local (Automatic) Data Block 10420 contains certain of the associated procedure's automatic data.

As described above, at each procedure call a MAS frame is constructed on top of the MAS of the domain in which the called procedure is executed. For example, when Procedure 10 calls Procedure 11 a Frame Header 10414 for Procedure 11 is constructed and placed on KOSMAS 10334. Procedure 11's linkage pointers are then generated, and placed in Procedure 11's Linkage Pointer 15 Block 10416. Next Procedure 11's local pointers are fhi generated and placed in Procedure 11's Local Pointer Block 10418. Finally, Procedure 11's local data is t placed in Procedure 11's Local Data Block 10420. During this operation, USERMAS 10328's frame label sequence is updated to include an entry pointing to Procedure 11's A Frame Header 10414. KOSMAS 10334's Stack Header 10410 is updated with respect to STO to the new top of KOSMAS 10334. Procedure 2's Frame Header 10414 is updated with respect to offset to Frame Header 10414 of Procedure 11 Frame 10412, and with respect to frame label sequence indicating location of Procedure 11's Frame Header 10414. As Procedure 11 is then the current procedure, FP is updated to a point between Linkage Pointer Block 10416 and Local Pointer Block 10418 of Procedure 11's Frame 10412. Also, as will be discussed below, a new frame is l -125- T Cl 1 1: :I kt.

1~I 41 constructed on SS 10336 or Procedure 11. CS 10110 will then proceed to execute Procedure 11. During execution of Procedure 11, any further local data generated may be placed on the top of Procedure 11's Local Data Block 10420. The top of stack offset information in Procedure 11's Frame Header 10414 and in KOSMAS 10334 Stack Header 10410 will be updated accordingly.

MAS's 10328 to 10334 thereby provide a per domain stack mechanism for storing data pertaining to individual procedures, thus allowing stacking of procedures without loss of this data. Although structured on a domain basis, MAS's 10328 to 10334 comprise a unified logical stack structure threaded together through information stored in MAS stack and frame headers.

15 As described above and previously, SS 10336 is a CS 10110 micromachine stack structure for storing, in part, CS 10110 micromachine state for each stacked VP 10310 procedure. Referring to Fig. 105, a partial schematic representation of a SS 10336 Stack Frame 10510 is shown.

SS 10336 Stack Header 10512 and Frame Headers 10514 contain information similar to that in MAS Stack Headers 10410 and Frame Headers 10414. Again, the information contained therein locates certain points within SS 10336 structure, and threads together SS 10336 with MAS's 10328 to 10334.

SS 10336 Stack Frame 10510 contains certain information used by the CS 10110 micromachine in executing the VP 10212 procedure with which this frame is -4'1 4 i -126associated. Procedure Pointer Block 10516 contains certain pointers including ABPs, used by CS 10110 micromachine in locating information within VP 10310's information structures. Micro-Routine Frames (MRFs) 10518 together comprise Micro-Routine Stack (MRS) 10520 within each SS 10336 Stack Frame 10510. MRS Stack 10520 is associated with the internal operation of CS 10110 microroutines executed during executionof the VP 10212 procedure associated with the Stack Frame 10510. SS 10336 is thus a dual function CS 10110 micromachine stack. Pointer Block 10516 entries effectively define an interface between CS 10110 micromachine and the current procedure of the current process. MRS 10520 comprise a stack mechanism for the internal operations of CS 10110 15 micromachine.

Having briefly described Virtual Processes 10212, SFURSM 10214 will be described next. As stated above, EURSM 10216 is similar in operation to FURSM 10214 and will be described in following detailed descriptions of CS 10110 structure and operation.

3. Referring again to Fig. 103, FURSM 10214 includes CS 10110 micromachine information structures used internally 8' 9. to CS 10110 micromachine in executing the procedures of a P 10310. When a VP, for example P 10310, is to be executed, certain information regarding that VP is transferred from the Virtual Processes 10212 to FURSM t f i -127- 10214 for use in executing that procedure. In this respect, FURSM 10214 may be regarded as an acceleration mechanism for the current Virtual Process 10212.

FURSM 10214 includes General Register File (GRF) 10354, Micro Stack Pointer Register Mechanism (MISPR) 10356, and Return Control Word Stack (RCWS) 10358. GRF 10354 includes Global Registers (GRs) 10360 and Stack Registers (SRs) 10362. GRs 10360 include Architectural Base Registers (ABRs) 10364 and Micro-Control Registers (MCRs) 10366. Stack Registers 10362 include Micro-Stack (MIS) 10368 and Monitor Stack (MOS) 10370.

Referring first to GRF 10354, and assuming for example that Procedure 11 of P 10310 is currently being executed, GRF 10354 primarily contains certain pointers S 15 to P 10310 data used in execution of Procedure 11. As previously disussed, CS 10110's addressing structure includes certain Architectural Base Pointers (ABP's) for each procedure. ABPs provide a framework for excessing CS 10110's address space. The ABPs of each procedure include a Frame Pointer a Procedure Base Pointer 4 4 *(PBP) and a Static Data Pointer (STP). As discussed above with reference to KOSPO 10318, these ABPs reside in the procedure's PEDs. When a procedure is called, these ABP's are transferred from that procedure's PED to ABR's 10364 and reside therein for the duration of that procedure. As indicated in Fig. 103, FP points between r Linkage Pointer Block 10416 and Local Pointer Blocks 10418 of Procedure 11's Frame 10412 on KOSMAS 10334. PBP points to the reference point from which the elements of KOSPO 10318 are located. SDP points to KOSSDA 10326. If -128- Procedure 11 calls, for example, a Procedure 12, Procedure 1l's ABPs will be transferred onto Procedure Pointer Block 10516 of SS 10336 Stack Frame 10510 for Procedure 11. Upon return to Procedure 11, Procedure 11's ABPs will be transferred from Procedure Pointer Block 10516 to ABR's 10364 and execution of Procedure 11 resumed.

MCRs 10336 contain certain pointers used by CS 10110 micromachine in executing Procedure 11.' CS 10110 micromachine pointers indicated in Fig. 103 include Program Counter Name Table Pointer (NTP), S- Interpreter Pointer (SIP), Secure Stack Pointer (SSP), and Secure Stack Top Offset (SSTO). NTP and SIP have been previously described with reference to KOSPO 10318 t, 15 and reside in KOSPO 10318. NTP and SIP are transferred into MCR's 10366 at start of execution of Procedure 11.

it PC, as indicated in Fig. 103, is a pointer to the Procedure 11 SIN currently being executed by CS 10110.

S PC is initially generated from Procedure 11's PBP and CEP and is thereafter incremented by CS 10110 micromachine as Procedure 11's SIN sequences are executed. SSP and SSTO are, as described in a following discussion, generated from information contained in SS 10336's Stack Header S, 10512 and Frame Headers 10514. As indicated in Fig. 103 25 SSP points to start of SS 10336 while SSTO indicates the current top frame on SS 10336, whether Procedure Pointer Block 10516 or a MRF 10518 of MRS 10520, by indicating an offset relative to SSP. If Procedure 11 calls a subsequent procedure, the contents of MCR's 10366 are r i -129transferred into Procedure 11's Procedure Pointer Block 10516 on SS 10336, and are returned to MCR's 10366 upon return to Procedure 11.

Registers 10360 contain further pointers, described in following detailed discussions of CS 10110 operation, and certain registers which may be used to contain the current procedure's local data.

Referring now to Stack Registers 10362, MIS 10368 is an upward extension, or acceleration, of MRS 10520 of the current procedure. As previously stated, MRS 10520 is used by CS 10110 micromachine in executing certain microroutines during execution of a particular procedure. MIS 10368 enhances the efficiency of CS 10110 cr,, micromachine in executing these microroutines by 15 accelerating certain most recent MRFs 10518 of that procedure's MRS 10520 into FU 10120. MIS 10368 may contain, for example, up to the eight most recent MRFs r 10518 of the current procedures MRS 10520. As various microroutines are called or returned from, MRS 10520 MRF's 10518 are transferred accordingly between SS 10336 and MIS 10368 so that MIS 10368 always contains at least the top MRF 10518 of MRS 10520, and at most.eight MRFs 10518 of MRS 10520. MISPR 10356 is a CS 10110 micromachine mechanism for maintaining MIS 10368. MISPR 10356 contains a Current Pointer, a Previous Pointer, and a Bottom Pointer. Current Pointer points to the top-most MRF 10518 on MIS 10368. Previous Pointer points to the i previous MRF 10518 on MIS 10368, and Bottom Pointer kA points to the bottom-most MRF 10518 on MIS 10368. MISPR 10356's Current, Previous and Bottom Pointers are updated a Y i

PP

I. aarm~ rrt ~-141 1111n11..

-130as MRFs 10518 are transferred between SS 10336 and MIS 10368. If Procedure 11 calls a subsequent procedure, all Procedure 11 MRFs 10518 are transferred from MIS 10368 to Procedure 11's MRS 10520 on SS 10336. Upon return to Procedure 11, up to seven of Procedure 11's MRFs 10518 frames are returned from SS 10336 to MIS 10368.

Referring to MOS 10370, MOS 10370 is a stack mechanism used by CS 10110 micromachine for certain microroutines for handling fault or error conditions.

These microroutines always run to completion, so that MOS 10370 resides entirely in FM 10120 and is not an extension of a stack residing in a P 10310 in MEM 10112.

MOS 10370 may contain, for example, eight frames. If more than eight successive fault or error conditions 15 occur, this is regarded as a major failure of CS 10110.

Control of CS 10110 may then be transferred to DP 10118.

As will be described in a following discussion, diagnostic programs in DP 10118 may then be used to diagnose and locate the CS 10110 faults or errors. In other embodiments of CS 10110 MOS 10370 may contain more or fewer stack frames, depending upon the degree of self diagnosis and correction capability desired for CS 10110.

RCWS 10358 is a two-part stack mechanism. A first part operates in parallel with MIS 10368 and a second part operates in parallel with MOS 10370. As previously described, CS 10110 is a microcode controlled system.

RCWS is a stack for storing the current microinstruction 'flu .e .9 .4 r* .4 4 4*f *49 *44tp .4 9 4r S*4 4

A"

-131being executed by CS 10110 micromachine when the current procedure is interrupted by a fault or error condition, or when a subsequent procedure is called. That portion of RCWS 10358 associated with MIS 10368 contains an entry for each MRF 10518 residing in MIS 10368. These RCWS 10358 entries are transferred between SS 10336 and MIS 10368 in parallel with their associated MRFs 10518. When resident in SS 10336, these RCWS 10358 entries are stored within their associated MRFs 10518. That portion of RCWS 10358 associated with MOS 10370 similarly operates in parallel with MOS 10370 and, like MOS 10370, is not an extension of an MEM 10112 resident stack.

In summary, each process active in CS 10110 exists as a separate, complete, and self-contained entity, or 15 Virtual Process, and is structurally organized on a domain basis. Each Virtual Process includes, besides procedure and data objects, a set of MAS's for storing local data of that processes procedures. Each Virtual Process also includes a CS 10110 micromachine stack, SS 10336, for storing CS 10110 micromachine state pertaining to each stacked procedure of the Virtual Process. CS ,2 10110 micromachine includes a set of information structures, register 10360, MIS 10368, MOS 10370, and RCWS 10358, used by CS 10110 micromachine in executing the Virtual Process's procedures. Certain of these CS 10110 micromachine information structures are shared with the currently executing Virtual Process, and thus are effectively acceleration mechanisms for the current Virtual Process, while others are completely internal to CS 10110 micromachine.

I 1 1 i -132- A primary feature of CS 10110 is that each process' macrostacks and secure stack resides in MEM 10112. CS 10110's macrostack and secure stacks are therefore effectively unlimited in depth.

Yet another feature of CS 10110 micromachine is the use of GRF 10354. GRF 10354. is, in an embodiment of CS 10110, a unitary register array containing for example, 256 registers. Certain portions, or address locations, of GRF 10354 are dedicated to, respectively, GRs 10360, MIS 10368, and MOS 10370. The capacities of GR 10360, MIS 10368, and MOS 10370, may therefore be adjusted, as required for optimum CS 10110 efficiency, by reassignment of GRF 10354's address space. In other embodiments of CS 10110, GRs 10360, MIS 10368, and MOS 10370 may be implemented a functionally separate registers arrays.

Having briefly described the structure and operation of Process Structures 10210, VP State Block 10218 will be S, described next below.

C. Vr tual -PPocees o---Satl-B-xsoc a Pr rocess- Great o n-~FiP-'.-l-9-2 Referring again to Fig. 102, VP State Blocks 10218 is used in management and control of processes. VP State Blocks 10218 contains a VP State Block for each Virtual Process (VP) selected for execution by CS 10110. Each such VP State Block contains at least the following information: the state, or identification number of a VP; -133entries identifying the particular principle and particular process of the VP; an AON pointer to that VP's secure stack SS 10336); the AON's of that VP's MAS stack objects MAS's 10328 to 10334); and, certain information used by CS 10110's VP Management System.

The information contained in each VP State Block thereby defines the current state of the asociated VP.

A Process is loaded into CS 10110 by building a primitive access record and loading this access record into CS 10110 to appear as an already existing VP. A VP is created by creating a Process Object, including pointers to macro-and secure-stack objects created for that VP, micromachine state entries, and a pointer to the user's program. CS 10110's KOS then generates Macro- and Secure-Stack Objects with headers for that process and, as described further below, loads protection information regarding that procesa objects into Protection Structures 10230. C IlU'10 KOS then copies this "I 4 primitive machine state record into a vacant VPSB selected by CS 10110's VP Manager, thus binding the newly S created VP into CS 10110. At that time a KOS Initializer procedure completes creation of the VP for example by calling in the user's program through a compiler. The newly creatd VP may then be executed by CS 10110.

to .Having briefly described VP State Blocks 10218 and creation of a VP, CS 10110's Addressing Structures 10220 will be decribed next below.

I 1 arnrr-~ PP ^L9-C-~ -134- D. Addressna-Struct-etu-eS-(-ie-3 r--O16-, 1. Obiects-.-UID-s-.-AGN -,-Namee andPihva-iede Addresses---Fi-:--tr6 As previously described, the data space accessible to CS 10110 is divided into segments, or containers, referred to as objects. In an embodiment of CS 10110, the addressable data space of each object has a capacity of 232 bits of information and is structured into 2 1 8 pages with each page containing 214bits of information.

Referring to Fig. 106A, a schematic representation of CS 10110's addressing structure is shown. Each object created for use in, or by operation of, a CS 10110 is permanently assigned a unique identifier (UID) An object's UID allows an object to be uniquely identified and located at any future point in time. Each UID is an 80 bit number, so that the total addressable space of all CS 10110's includes 2 objects wherein each object may contain up to 2 bits of information. As indicated in Fig. 106, each 80 bit UID is comprised of 32 bits of Logical Allocation Unit Identifier (LAUID) and 48 bits of Object Serial Number (OSN). LAUIDs are associated with individual CS 10110 systems. LAUIDs identify the particular CS 10110 system generating a particular object. Each LAUID is comprised of a Logical Alloc$tion Unit Group Number (LAUGN) and a Logical Allocation Unit Serial Number (LAUSN). LAUGNs are assigned to individual CS 10110 systems and may be guaranteed to be unique to a f. Sf w.

-135- *L 0 r a S :oc particular system. A particular system may, however, be assigned more than one LAUGN so that there may be a time varying mapping between LAUGNs and CS 10110 systems.

LAUSNs are assigned within a particular system and, while LAUSNs may be unique within a particular system, LAUSNs need not be unique between systems and need not map onto the physical structure of a particular system.

OSNs are associated with individual objects created by an LAU and are generated by an Architectural Clock in each CS 10110. Architectural clock is defined as a 64 bit binary number representing increasing time. Least significant bit of architectural clock represents increments of 600 picoseconds, and most significant bit represents increments of 127 years. In the present 15 embodiment of CS 10110, certain most significant and least significant bits of architectural clock time are disregarded as generally not required practice. Time indicated by architectural clock is measured relative to an arbitrary, fixed point in time. This point in time is the same for all CS 10110s which will ever be constructed. All CS 10110s in existence will therefore indicate the same architectural clock time and all UIDs generated will have a common basis. The use of an architectural clock for generation of OSNs is 25 advantageous in that it avoids the possibility of accidental duplication of OSNs if a CSC 10110 fails and is subsequently reinitiated.

I

j'U o 4 44 0 0000 44 4* 4

ES

4, 04 4 41 -136- As stated above, each object generated by or for use in a CS 10110 is uniquely identified by its associated UID. By appending Offset and Length information to an object's UID, a UID logical address is generated which may be used to locate particular segments of data residing in a particular object. As indicted in Fig.

106, O and L fields of a UID logical address are each 32 bits. O and L fields can therefore indicate any 32-1 particular bit, out of 2. bits, in an object and thus lo allow bit granular addressing of information in objects.

As indicated in Fig. 106 and as previously described, each object active in CS 10110 is assigned a short temporary unique identifier valid only within JP 10114 referred to as an Active Object Number (AON).

Because fewer objects may be active in a CS 10110 than may exist in a CS 10110's address space, AON's are, in the present embodiment of CS 10110, 14 bits in length. A particular CS 10110 may therefore contain up to 214 active objects. An object's AON is used within JP 10114 in place of that object's UID. For example, as discussed j F above with reference to process structures 10210, a SI procedure's FP points to start of that procedure's frame on its process' MAS. When that FP is residing in SS 10336, it is expressed as a UID. When that procedure is to be executed, FP is transferred from SS 10336 to ABR's 10364 and is translated into the corresponding AON. i Similarly, when that procedure is stacked, FP is returned to SS 10336 and in doing so is translated into the corresponding UID. Again, a particular data segment in i. .I -137an object may be addressed by means of an AON logical address comprising the object's AON plus associated 32 bit Offset and Length fields.

Each operand appearing in a process is assigned a Name and all references to a process's operands are through those assigned Names. As indicated in Fig. 106B, in the present embodiment of CS 10110 each Name is an 8, 12, or 16 bit number. All Names within.a particular process will be of the same length. As will be described in a following.discussion, Names appearing during execution of a process may be resolved, through a procedure's Name Table 10350 or through Name Cache 10226, to an AON logical address. As described below, an AON logical address corresponding to an operand Name may then 1i be evaluated to a MEM 10112 physical address to locate the operand referred to.

.O The evaluation of AON logical addresses to MEM 10112 t t I' physical addresses is represented in Fig. 106C. An AON logical address's L field is not involved in evaluation of an AON logical address to a physical address and, for purposes of clarity of presentation, is therefore not represented in Fig. 106C. AON logical address L field is to be understood to be appended to the addresses txt represented in the various steps of the evaluation procedure shown in Fig. 106C.

As described above, objects are 2 32 bits structured into 2 18 pages with each page containing a 2 14 bits of j r i -138data. MEM 10112 is similarly physically structured into frames with, in the present embodiment of CS 10110, each frame containing 214 bits of data. In other embodiments of CS 10110, both pages and frames may be of different sizes but the translation of AON logical addresses to MEM 10112 physical addresses will be similar to that described momentarily.

An AON logical address 0 field was previously described as a 32 bit number representing the start, relative to start of the object, of the addressed data segment within the object. The 18 most significant bits of 0 field represent the number of the page within the object upon which the first bit of the addressed data occurs. The 14 least significant bits of 0 field represent the offset (0 relative to the start of the page, within that page of the first bit of the addressed data. AON logical address 0 field may therefore, as indicated in Fig. 106C, be divided into an 18 bit page field and a 14 bit offset within page (Op) field.

Since, as described above, MEM 10112 physical frame size is equal to object page size, AON logical address Op field may be used directly as an offset within frame (Of) field of the physical address. As will be described below, an AON logical address AON and P fields may then 25 be related to the frame number (FN) of the MEM 10112 frame in which that page resides, through Addressing Mechanisms 10220.

t tt t C .4, *4 I 4 4, $1 I 14 "I i -139- Having briefly described the relationships between UIDs, UID Logical Addresses, Names, AONs, AON Logical Addresses, and MEM 10112 Physical Addresses, Addressing Mechanisms 10220 will be described next below.

2. Address1nc0-Mechean2sms- e220 (Pig. 107 Referring to Fig. 107, a schematic representation of Computer System 10110's Addressing Mechanisms 10220 is shown. As previously described, Addressing Mechanisms 10220 comprise UID/AON Tables 10222, Memory Management Tables 10224, Name Cache 10226, and Address Translation Unit 10228.

UID/AON Tables 10222 relate each object's UID to its assigned AON and include AOT Hash Table (AOTHT) 10710, t 1 Active Object Table (AOT) 10712, and Active Object Table Annex (AOTA) 10714.

SAn AON corresponding to a particular UID is determined through AOTHT 10710. The UID is hashed to I provide a UID index into AOTHT 10710, which then provides *the corresponding AON. AOTHT 10710 is effectively an acceleration mechanism of AOT 10712 to, as just S* described, provide rapid translation of UIDs to AONs.

r 25 AONs are used as indexes into AOT 10712, which provides a corresponding AOT Entry (AOTE). An AOTE as described in following detailed discussions of CS 10110, includes, among other information, the UID corresponding to the AON indexing the AOTE. In addition to providing translation i 9i -140between AONs and UIDs, the UID of an AOTE may be compared to an original UID to determine the correctness of an AON from AOTHT 10710.

Associated with AOT 10712 is AOTA 10714. AOTA 10714 is an extension of AOT 10712 and contains certain information pertaining to active objects, for example the domain of execution of each active procedure object.

Having briefly described CS 10110's mechanism for relating UIDs and AONs, CS 10110's mechanism for 10 resolving operand Names to AON logical addresses will be described next below.

SC 3. Name- -Resolution "r Referring first to Fig. 103, each procedure object in S< a VP, for example KOSPO 10318 in VP 10310, was described as containing a Name Table (NT) 10350. Each NT 10350 j contains a Name Table Entry (NTE) for each operand whose So' jName appears its procedure. Each NTE contains a I .description of how to resolve the corresponding Name to .an AON Logical Address, including fetch mode information, 20 type of data referred to by that Name, and length of the data segment referred to.

Referring to Fig. 108, a representation of an NTE is shown. As indicated, this NTE contains seven information fields: Flag, Base Predisplacbment Length I Displacement Index and Inter-element Spacing (IES). Flag Field, in part, contains information describing how the remaining fields of the NTE are to be interpreted, type of information referred to by the NTE, and how that information is to handled when fetched from i *04 food 0 0 *4 of 4,4 0 *i S 4 4 4 54r 4 4 4*4 44 4 4 .4i 4 4 *4i -141- MEM 10112. L Field, as previously described, indicates length, or number of bits in, the data segment.

Functions of the other NTE fields will be described during the following discussions.

In a present embodiment of CS 10110, there are five types of NTE: base is not a Name, address resolution is not indirect; B is not a Name, address resolution is indirect; B is a Name, address resolution is indirect; B is a Name, address resolution is indirect. A fifth type is an NTE selecting a particular element from an array of elements. These five types of NTE and their resolution will be described below, in the order mentioned.

In the first type, B is not a Name and address 15 resolution is not indirect, B Field specifies an ABR 10364 containing an AON plus offset (AON/0) Pointer. The contents of D Field are added to the 0 Field of this pointer, and the result is the AON Logical Address of the operand. In the second type, B is not a Name and address resolution is indirect, B Field again specifies an ABR 10364 containing an AON/O pointer. The contents of PR Field are added to the 0 Field of the AON/0 pointer to provide an AON Logical Address of a Base Pointer. The Base Pointer AON Logical Address is evaluated, as described below, and the Base Pointer fetched from MEM 10112. The contents of D Field are added to the 0 Field of the Base Pointer and the result is the AON Logical Address of the operand.

C1 -142- NTE types 3 and 4 correspond, respectively to NTE types 1 and 2 and are resolved in the same manner except that B Field contains a Name. The B Field Name is resolved through another NTE to obtain an AON/O pointer which is used in place of the ABR 10364 pointers referred to in discussion of types 1 and 2.

The fifth type of NTE is used in references to elements of an array. These array NTEs are resolved in the same manner as NTE types 1 through 4 above to provide 10 an AON Logical Address of the start of the array. I and IES Fields provide additional information to locate a particular element in the array. I Field is always Name which is resolved to obtain an operand value representing Stt' the particular element in the array. IES Field provides information regarding spacing between elements of the array, that is the number of bits between adjacent Selement of the array. IES Field may contain the actual IES value, or it may contain a Name which is resolved to an AON Logical Address leading to the inter-element spacing value. The I and IES values, obtained by resolving the I and IES Fields as just described, are multiplied together to determine the offset, relative to I the start of the array, of the particular element referred to by the NTE. This within array offset is added to the 0 Field of the AON Logical Address of the start of the array to provide the AON Logical Address of the element.

t 0 -143- In the current embodiment of CS 10110, certain NTE fields, for example B, D, and Flag fields, always contain literals. Certain other fields, for example, IES, D, PRE, and L fields, may contain either literals or names to be resolved. Yet other fields, for example I field, always contain names which must be resolved.

Passing of arguments from a calling procedure to a called procedure has been previously discussed with reference to Virtual Processes 10212 above, and more 10 specifically with regard to MAS's 10328 to 10334 of VP 10310. Passing of arguments is accomplished through the

S

t calling and called procedure's Name Tables 10350. In Sillustration, a procedure W(a,b,c) may wish to pass arguments a, b, and c to procedure where arguments, v and w correspond to arguments a, b, and c.

At compilation, NTEs are generated for arguments a, b, c and c in Procedure W's procedure object, and NTEs are t t t generated for arguments u, v and w in Procedure X's S. procedure object. Procedure X's NTEs for u, v, and w are 20 constructed to resolve to point to pointers in Linkage Pointer Block 10416 of Procedure X's Frame 10412 in MAS.

To pass arguments a, b, and c from Procedure W to Procedure X, the NTEs of arguments a, b, and c are Sresolved t AON Logical Addresses AON/O form).

Arguments a, b, and c's AON Logical Addresses are then translated to corresponding UID addresses which are placed in Procedure X's Linkage Pointer Block 10416 at those places pointed to by Procedure X's NTEs for u, v, and w. When Procedure X is executed, the resolution of t -144- Procedure X's NTEs for u, v, and w will be resolved to locate the pointers, in Procedure X's Linkage Pointer Block 10416 to arguments a, b, and c. When arguments are passed in this manner, the data type and length information are obtained from the called procedure's NTEs, rather than the calling procedure's NTEs. This allows the calling procedure to pass only a portion of, for example, arguments a, b, or c, to the called procedure and thus may be regarded as a feature of CS 10 10110's protection mechanisms.

Having briefly described resolution of Names to AON/Offset addresses, and having previously described Stranslation of UID addresses to AON addresses, the "Ii, *evaluation of AON addresses to MEM 10112 physical addresses will be described next below.

4. Evalua-tion- of- AON--Addresses--to--Phygica Addresses Fig~-i-- H}-7 Referring again to Fig. 1 0 7 f a partial schematic *o representation of CS 10110's Memory Management Table 10224 is shown. Memory Hash Table (MHT) 10716 and Memory Frame Table (MFT) 10718 are concerned with translation of SA. ON addresses into MEM 10112 physical addresses and will be discussed first. Working Set Matrix (WSM) 10720 and Virtual Memory Manager Request Queue (VMMRQ) 10722 are concerned with management of MEM 10112'S available physical address base and will be discussed second.

Active Object Request Queue (AORQ) 10728 and Logical -145- Allocation Unit Directory (LAUD) 10730 are concerned with locating inactive objects and management of which objects are active in CS 10110 and will be discussed last.

Translation of AON/O Logical Addresses to MEM 10112 physical addresses was previously discussed with reference to Fig. 106C. As stated in that discussion, objects are divided into pages. Correspondingly, the AON/O Logical Address' 0 Field is divided into an 18 bit page number Field and a 14 bit offset within a page (Op) Field. MEM 10112 is structured into frames, each of which in the present embodiment of CS 10110 is equal to a page of an object. An AON/O address' 0 Field may therefore be used directly as an offset within frame (Of) Sof the corresponding physical address. The AON and P fields of an AON address must, however, be translated into a MEM 10112 frame represented by a corresponding I fFrame Number (FN).

t t Referring now to Fig. 107, an AON address' AON and P Fields are "hashed" to generate an MHT index which is used as an index into MHT 10716. Briefly, "hashing" is a method of indexing, or locating, information in a table wherein indexes to the information are generated from the information itself through a "hashing function". A hashing function maps each piece of information to the corresponding index generated from it through the hashing function. MHT 10716 then provides the corresponding FN of the MEM 10112 frame in which that page is stored. FNs are used as indexes into MFG 10718, which contains, for each FN, an entry describing the page stored in that frame. This information includes the AON and P of the -146page stored in that MEM 10112 frame. An FN from MHT 10716 may therefore be used as an index into MFT 10718 and the resulting AON/P of MFT 10718 compared to the original AON/P to confirm the correctness of the FN obtained from MHT 10716. MHT 10716 is an effectively acceleration mechanism of MFT 10718 to provide rapid translation of AON address to MEM 10112 physical addresses.

j MFT 10718 also stores "used" and "modified" 10 information for each page in MEM 10112. This information j 0 indicates which page frames stored therein have been used i and which have been modified. This information is used by CS 10110 in determining which frames may be deleted from MEM 10112, or are free, when pages are to be written into MEM 10112 from backing store (ED 10124) For Stexample, if a page's modified bit indicates that that S page has not been written into, it is not necessary to write that page back into backing store when it is deleted from MEM 10112; instead, that page may be simply S' 20 erased.

Referring finally to ATU 10228, ATU 10228 is an S acceleration mechanism for MHT 10716. AON/O addresse, are used directly, without hashing, as indexes into ATU S 10228 and ATU 10228 correctly provides corresponding FN and 0 outputs. A CS 10110 mechanism, described in a following detailed discussion of CS 10110 operation -147continually updates the contents of ATU 10228 so that ATU 10228 contain the FN's and Op (Of of the pages most frequently referenced by the current process. If ATU 10228 does not contain a corresponding entry for a given AON input, an ATU fault occurs and the FN and 0 information may be obtained directly from MHT 10716.

Referring now to WSM 10720 and VMMRQ 10722, as previously stated these mechanisms are concerned with the S. management of MEM 10112's available address space. For example, if MHT 10716 and MFT 10718 do not contain an entry for a page referenced by the current procedure, an w s' MHT/MFT fault occurs and the reference page must be I .fetched from backing store (ED 10124) and read into MEM 10112. WSM 10720 contains an entry for each page resident in MEM 10112. These entries are accessed by indexes comprising the Virtual Processor Number (VPN) of the virtual process making a page reference and the P of 1 the ,age being referenced. Each WSM 10720 entry contains S2 bits stating whether the particular page is part of a 20 VP's working set, that is, used by that VP, and whether that page has been referenced by that VP. This information, together with the information contained in that MFT 10718 entries described above, is used by CS 10110's Virtual Memory Manager (VMM) in transferring pages into and out of MEM 10112.

CS 10110's VMM maintains VMMRQ 10722, which is used by VMM to control transfer of pages into and out of MEM 10112. VMMRQ 10722 includes Virtual Memory Request Counter (VMRC) 10724 and a Queue of Virtual Memory h i d A 1 -148- Request Entries (VMREs) 10726. As will be discussed momentarily, VMRC 10724 tracks the number of currently outstanding request for pages. Each VMRE 10726 describes a particular page which has been requested. Upon occurrence of a MHT/MFT (or page) fault, VMRC 10724 is incremented, which initiates operation of CS 10110's VMM, and a VMRE 10726 is placed in the queue. Each VMRE 10726 comprises the VPN of the process requesting the page and the AON/O of the page requested. At this time, the VP making the request is swapped out of JP 10114 and another VP bound to JP 10114. VMM allocates MEM 10112 frame to Scontain the requested page, using the previously 6IP° described information in MFT 10718 and WSM 10720 to select this frame. In doing so, VMM may discard a page currently resident in MEM 10112 for example, on the basis of being the oldest page, an unused page, or an unmodified page which does not have to be written back Sj into backing store. VMM then requests an I/O operation to transfer the requested page into the frame selected by S" 20 the VMM. While the I/O operation is proceeding, VMM generates new entries in MHT 10716 and MFT 10718 for the .requested page, cleans the frame in MEM 10112 which is to J be occupied by that page, and suspends operation. IOS *101,16 will proceed to execute the I/O operation and writes the requested page directly into MEM 10112 in the frame specified by VMM. IOS 10116 then notifies CS f 10110's VMM that the page now resides in memory and can be referenced. At some later time, that VP requesting that page will resume execution and repeat that L 1 1 o I r l r- -149- 4099 0 0, tt' I~r~ ~cc I II *4 4 4*40 till..

t 4 reference. Going first to ATU 10228, that VP will take an ATU 10228 fault since VP 10212 has not yet been updated to contain that page. The VP will then go to MHT 10716 and MFT 10718 for the required information and, concurrently, WSM 10720 and ATU 10228 will be updated.

In regard to the above operations, each VP active in CS 10110 is assigned a Page Fault Frequency Time Factor (PFFT) which is used by CS 10110's VMM to adjust that VP's working set so that the interval between-successive page faults for that VP lies in an optimum time range.

This assists in ensuring CS 10110's VMM is operating most efficiently and allows CS 10110's VMM to be tuned as required.

The above discussions have assumed that the page being referenced, whether from a UID/O address, an AON/O address, or a Name, is resident in an object active in CS 10110. While an object need not have a page in MEM 10112 to be active, the object must be active to have a page in MEM 10112. A VP, however, may reference a page in an 20 object not active in CS 10110. If such a reference is made, the object must be made active in CS 10110 before the page can be brought into MEM 10112. The result is an operation similar to the page fault operation described above. CS 10110 maintains an Active Object Manager (AOM), including Active Object Request Queue (AORQ) 10728, which are similar in operation to CS 10110's VMM and VMMRQ 10722. CS 10110's AOM and AORQ 10728 operate in conjunction with AOTHT 10710 and AOT 10712 to locate inactive objects and make them active by assigning them -150- AON's and generating entries for them in AOTHT 10710, AOT 10712, and AOTA 10714.

Before a particular object can be made active in CS 10110, it must first be located in backing store (ED 10124). All objects on backing store are located through a Logical Allocation Unit Directory (LAUD) 10730, which is resident in backing store. An LAUD 10730 contains entries for each object accessible to the particular CS 10110. Each LAUD 10730 entry contains the information 10 necessary to generate an AOT 10712 entry for that object. An LAUD 10730 is accessed through a UID/O address contained in CS 10110's VMM. A reference to an LAUD 10730 results in MEM 10112 frames being assigned to that LAUD 10730, and LAUD 10730 being transferred into MEM 10112. If an LAUD 10730 entry exists for the referenced inactive object, the LAUD 10730 entry is "e transferred into AOT 10712. At the next reference to a t page in that object, AOT 10712 will provide the AON for that object but, because the page has not yet been S 20 transferred into MEM 10112, a page fault will occur.

This page fault will be handled in the manner described above and the referenced page transferred into MEM 10112.

Having briefly described the structure and operation of CS 10110's Addressing Structure, including the relationship between UIDs, Names, AONs, and Physical Addresses and the mechanisms by which CS 10110 manages

/%I

-151the available address space of MEM 10112, CS 10110's protection structures will be described next below.

E. CS 10O1ll0'f -Potect-ion--Mechanisms--g4q--g99 Referring to Fig. 109, a schematic representation of Protection Mechanisms 10230 is shown. Protection Tables 10232 include Active Primitive Access Matrix (APAM) 10910, Active Subject Number Hash Table (ASNHT) 10912, and Active Subject Table (AST) 10914. Those portions of Protection Mechanism 10230 resident in FU 10120 include 10 ASN Register 10916 and Protection Cache (PC) 10234.

SAs previously discussed, access rights to objects are arbitrated on the basis of subjects. A subject has been defined as a particular combination of a Principle, Process, and Domain (PPD), each of which is identified by a corresponding UID. Each object has associated with it an Access Control List (ACL) 10918 containing an ACL Entry (ACLE) for each subject having access rights to that object.

When an object becomes active in CS 10110 is 20 assigned an AON) each ACLE in that object's ACL 10918 is written into APAM 10910. Concurrently, each subject having access rights to that object, and for which there is an ACLE in that object's ACL 10918, is assigned an r Active Subject Number (ASN). These ASNs are written into ASNHT 10912 and their corresponding PPDs are written into AST 10914. Subsequently, the ASN of any subject requesting access to that object is obtained by hashing 1

I

-152the PPD of that subject to obtain a PPD index into ASNHT 10912. ASNHT 10912 will in turn provide a corresponding ASN.' An ASN may be used as an index into AST 10914. AST 10914 will provide the corresponding PPD, which may be I 5 compared to an original PPD to confirm the accuracy of the ASN.

As described above, APAM 10910 contains an ACL 10918 for each object active in CS 10110. The access rights of any particular active subject to a particular active 10 object are determined by using that subject's ASN and that object's AON as indexes into APAM 10910. APAM 1091Q in turn provides a 4 bit output defining whether that subject has Read Write or Execute rights with j I respect to that object, and whether that particular entry 15 is Valid ASN Register 10916 and PC 10234 are effectively t t" acceleration mechanisms of Protection Tables 10232. ASN Register 10916 stores the ASN of a currently active subject while PC 10234 stores certain access right S. 20 information for objects being used by the current process. PC 10234 entries are indexed by ASNs from ASN register 10916 and by a mode input from JP 10114. Mode o. input defines whether the current procedure intends to i n read, write, or execute with respect to a particular object having an entry in PC 10234. Upon receiving ASN and mode inputs, PC 10234 provides a go/nogo output indicating whether that subject has the access rights i

S

t' 1 -153required to execute the intended operation with respect to that object.

In addition to the above mechanism, each procedure to which arguments may be passed in a cross-domain call has associated with it an Access Information Array (AIA) 10352, as discussed with reference to Virtual Processes 10212. A procedure's AIA 10352 states what access rights a calling procedure (subject) must have to a particular object (argument) before-the called procedure can operate i^ 10 on the passed argument. CS 10110's protection mechanisms 1 Iwo compare the calling procedures access rights to the rights required by the called procedure. This insures too,$# the calling procedure may not ask a called procedure to do what the calling procedure is not allowed to do.

Effectively, a calling procedure can pass to a called procedure only the access rights held by the calling C'e' procedure.

Finally, PC 10234, APAM 10910, or AST 10914 faults misses) are handled in the same manner as described above with reference 'to page faults in discussion of CS 10110's Addressing Mechanisms 10220. As such, the handling of protection misses will not be 4' discussed further at this point.

r :Having briefly described structure and operation of CS 10110's Protection Mechanisms 10230, CS 10110's Micro- Instruction Mechanisms 10236 will be described next below. r 1 -154- F. CS-10110--M cro-nstrcet:on-Mechansms (Fig. 110) As previously described, CS 10110 is a multiple language machine. Each program written in a high level user language is compiled into a corresponding S-Language program containing instructions expressed as SINs. CS 10110 provides a set, or dialect, of microcode instructions, referred to as S-Interpreters (SINTs) for each S-Language. SINTs interpret SINs and provide corresponding sequences of microinstructions for detailed o* 10 control of CS 10110.

Referring to Fig. 110, a partial schematic representation of CS 10110's Micro-Instruction Mechanisms 10236 is shown. At system initialization all CS 10110 microcode, including SINTs and all machine assist microcode, is transferred from backing store to Micro- Code Control Store (mCCS) 10238 in MEM 10112. The Micro- S. Code is then transferred from mCCS 10238 to FU Micro-Code Structure (FUmC) 10240 and EU Micro-Code Structure (EUmC) 10242. EUmC 10242 is similar in structure and operation to FUmC 10240 and thus will be described in following detailed descriptions of CS 10110's structure and operation. Similarly, CS 10110 machine assist microcode Swill be described in following detailed discussions. The I present discussion will concern CS 10110's S-Inherpreter mechanisms.

CS 10110's S-Interpreters (SINTs) are loaded into S- Interpret Table (SITT) 11012, which is represented in Fig. 110 as containing S-Interpreters 1 to N. Each SIT t -155contains one or more sequences of micro-code; each sequence of microcode corresponds to a particular SIN in that S-Language dialect. S-Interpreter Dispatch Table (SDT) 11010 contains S-Interpreter Dispatchers (SDs) 1 to N. There is one SD for each SINT in SITT 11012, and thus a SD for each S-Language dialect. Each SD comprises a set of pointers. Each pointer in a particular SD corresponds to a particular SIN of that SD's dialect and points to the corresponding sequence of microinstructions 10 for interpreting that SIN in that dialect's SIT in SITT 11012. In illustration, as previously discussed when a «particular procedure is being executed the SIP for that procedure is transferred into one of mCR's 10366. That SIP points to the start of the SD for the SIT which is to be used to interpret the SINs of that procedure. In Fig.

110, the SIP in mCRs 10366 is shown as pointing to the start of SD2. Each S-Op appearing during execution of t t .that procedure is an offset, relative to the start of the selected SD, pointing to a corresponding SD pointer.

That SD pointer in turn points to the corresponding sequence of microinstructions for interpreting that SIN in the corresponding SIT in SITT 11012. As will be t "described in following discussions, onc the start of a t microcode sequence for interpreting an SIN has been selected, CS 10110 micromachine then proceeds to sequentially call the microinstructions of that sequence from SITT 11012 and use those microinstructions to control operation of CS 10110. t. t -156- G. Summary of--Certain-CS-ltOH0-Peatt es-and Alternate- Embodiments The above Introductory Overview has described the overall structure and operation and certain features of CS 101, that is, CS 10110. The above Introduction has further described the structure and operation and further features of CS 10110 and, in particular, the physical implementation and operation of CS 10110's information, control, and addressing mechanisms. Certain of these CS 10110 features are summarized next below to briefly state S*the basic concepts of these features as implemented in CS L 10110. In addition, possible alternate embodiments of CatI certain of these concepts are described.

First, CS 10110 is comprised of a plurality of independently operating processors, each processor having a separate microinstruction control. In the present c embodiment of CS 10110, these processors include FU 10120, EU 10122, MEM 10112 and IOS 10116. Other such independently operating processors, for example, special 20 arithmetic processors such as an array processor, or multiple FU 10120's, may be added to the present CS 10110.

S" In this regard, MEM 10112 is a multiport processor S ,*having one or more separate and independent ports to each processor in CS 10110. All communications between CS 10110's processors are through MEM 10112, so that MEM 10112 operates as the central communications node of CS l i 7 I, hL -157- 10110, as well as performing memory operations. Further separate and independent ports may be added to MEM 10112 as further processors are added to CS 10110. CS 10110 may therefore be described as comprised of a plurality of separate, independent processors, each having a separate microinstruction control and having a separate and independent port to a central communications and memory node which in itself is an independent processor having a separate and independent microinstruction control. As will be further described in a following detailed description of MEM 10112, MEM 10112 itself is comprised of a plurality of independently operating processors, each performing memory related operations and each having a separate microinstruction control. Coordination of operations between CS 10110's processors is achieved by passing "messages" between the processors, for example, SOP's and descriptors.

CS 10110's addressing mechanisms are based, first, 2 upon UID addressing of objects. That is, all information S 20 generated for use in or by operation of a CS 10110, for example, data and procedures, is structured into objects S and each object is assigned a permanent UID. Each UID is s t unique within a particular CS 10110 and between all CS S' 10110's and is permanently associated with a particular object. The use of UID addressing provides a permanent, unique addressing means which is common to all CS 10110's, and to other computer systems using CS 10110's UID addressing.

l -158- 4+ 44 5* i *4 As 4i *iS Effectively, UID addressing means that the address (or memory) space of a particular CS 10110 includes the address space of all systems, for example disc drives or other CS 10110s, to which that particular CS 10110 has access. UID addressing allows any process in any CS 10110 to obtain access to any object in any CS 10110 to which it has physical access, for example, another CS 10110 on the other side of the world. This access is constrained only by CS 10110's protection mechanism. In 10 alternate embodiments of CS 10110, certain UIDs may be set aside for use only within a particular CS 10110 and may be unique only within that particular CS 10110.

These reserved UIDs would, however, be a limited group known to all CS 10110 systems as not having uniqueness 15 between systems, so that the unique object addressing capability of CS 10110's UID addressing is preserved.

As previously stated, AONs and physical descriptors are presently used for addressing within a CS 10110, effectively as shortened UIDs. In alternate embodiments of CS 10110, other forms of AONs may be used, or AONs may be discarded entirely and UIDs used for addressing within as well as between CS 10110s.

CS 10110's addressing mechanisms are also based upon the use of descriptors within and between CS 10110s.

Each descriptor includes an AON or UID field to identify a particular object, an offset field to specify a bit granular offset within the object, and a length field to specify a particular number of bits beginning at the specified offset. Descriptors may also include a type, or format field identifying the particular format of the 5555 a. S

I

t ?L~b~ r

'E

B

QI

,i 4K -159data referred to by the descriptor. Physical descriptors are used for addressing MEM 10112 and, in this case, the AON or UID field is replaced by a frame number field referring to a physical location in MEM 10112.

As stated above, descriptors are used for addressing within and between the separate, independent processors (FU 10120, EU 10122, MEM 10112, and IOS 10116) comprising CS 10110, thereby providing common, system wide bit granular addressing which includes format information.

In particular, MEM 10112 responds to the type information fields of descriptors by performing formatting operations •to provide requestors with data in the format specified 46; .44 by the requestor in the descriptor. MEM 10112 also accepts data in a format specified in a descriptor and reformats that data into a format most efficiently used by MEM 10112 to store the data.

As previously described, all operands are referred to t* in CS 10110 by Names wherein all Names within a particular S-Language dialect are of a uniform, fixed a* a 20 size and format. A K value specifying Name size is provided to FU 10120, at each change in S-Language dialect, and is used by FU 10120 in parsing Names from L ;the instruction stream. In an alternate embodiment of CS 10110, all Names are the same size in all S-Language dialects, so that K values, and the associated circuitry in FU 10120's parser, are not required.

L i -C ^c -160- 4i 44..

i I I Ct i Slt I Si 4 Ir 9 S ei 4 9.4.4i 4 Finally, in descriptions of CS 10110's use of SOPs, FU 10120's microinstruction circuitry was described as storing one or more S-Interpreters. S-Interpreters are sets of sequences of microinstructions for interpreting the SOPs of various S-Language dialects and providing corresponding sequences of microinstructions to control CS 10110. In an alternate embodiment of CS 10110, these S-Interpreters (SITT 11012) would be stored in MEM 10112. FU 10120 would receive SOPs from the instruction 10 stream and, using one or more S-Interpreter Base Pointers (that is, architectural base pointers pointing to the SITT 11012 in MEk 10112), address the SITT 11012 stored in MEM 10112. MEM 10112 would respond by providing, from the SITT 11012 in MEM 10112, sequences of 15 microinstructions to be used directly in controlling CS 10110. Alternately, the SITT 11012 in MEM 10112 could provide conventional instructions usable by a conventional CPU, for example, Fortran or machine language instructions. This, for example, would allow FU 10120 to be replaced by a conventional CPU, such as a Data General Corporation Eclipse*.

Having briefly summarized certain features of CS 10110, and alternate embodiments of certain of these features, the structure and operation of CS 10110 will be described in detail below.

p "i ct -161- 2. DETAiLtED- DESCRgPTIOOP-ES- O_-Mf-JeR-SUBSSTfMS Having previously described the overall structure and operation of CS 10110, the structure and operation of CS 10110's major subsystems will next be individually described in further detail. As previously discussed, CS 10110's major subsystems are, in the order in which they will be described, MEM 10112, FU 10120, EU 10122, 103 10116, and DP 10118. Individual block diagrams of MEM 1 0 10112, FU 10120, EU 10122, IOS 10116, and DP 10118 are I shown in, respectively, Figures 201 through 205. Figures t 201 through 205 may be assembled as shown in Fig. 206 to r construct a more detailed block diagram of CS 10110 corresponding to that shown in Fig. 101. For the purposes of the following descriptions, it is assumed that Figs. 201 through 205 have been assembled as shown in Fig. 206 to construct such a block diagram. Further diagrams will be presented in following descriptions as required to convey structure and operation of CS 10110 to one of ordinary skill in the art.

As previously described, MEM 10112 is an intelligent, priortizing memory having separate and independent ports MIO 10128 and MJP 10140 to, respectively, IOS 10116 and JP 10114. MEM 10112 is shared by and is accessible to both JP 10114 and IOS 10116 and is the primary memory of CS 10110. In addition, MEM 10112 is the primary path for information transferred between the external world (through IOS 10116) and JP 10114.

L h* i -162- As will be described further below, MEM 10112 is a two-level memory providing fast access to data stored therein. MEM 10112 first level is comprised of a large set of random access arrays and MEM 10112 second level is comprised of a high speed cache whose operation is generally transparent to memory users, that is JP 10114 and IOS 10116. Information stored in MEM 10112, in either level, appears to be bit addressable to both JP 10114 and IOS 10116. In addition, MEM 10112 presents 10 simple interfaces to both JP 10114 and IOS 10116. Due to a high degree of pipe lining (concurrent and overlapping memory operations) MEM 10112 interfaces to both JP 10114 and IOS 10116 appear as if each HP 10114 and IOS 10116 have full access to MEM 10112. This feature allows data transfer rates of up to, for example, 63.6 megabytes per second from MEM 10112 and 50 megabytes per second to MEM 10112.

In the following descriptions, certain terminology c 9e used on those descriptions will be introduced first, followed by description of MEM 10112 physical organization. Then MEM 10112 port structures will be Ctt,. described, followed by descriptions of MEM 10112's control organization and control flow. Next, MEM 10112's interfaces to JP 10114 and IOS 10116 will be described.

Following these overall descriptions the major logical structures of MEM 10112 will be individually described, starting at MEM 10112's interfaces to JP 10114 and IOS

.J

1 r I I- 'it 4 9f *t4t OC C c fCC 44 4 I. I it 4* II i I -163- 10116 and proceeding inwardly to MEM 10112's first (or bulk) level of data stored. Finally, certain features of MEM 10112 microcode control structure will be described.

A. MEM 10-1-2---Pfs-:-20- 206T-20-7 a. Termino-loy Certain terms are used throughout the following descriptions and are defined here below for reference by the reader.

A word is 32 bits of data 10 A byte is 8 bits of data A block is 128 bits of data (that is, 4 words).

A block is always aligned on a block boundary, that is the low order 7 bits of logical or physical address are zero (see Chapter 1, Sections A.f and D. Descriptions of 15 CS 10110 Addressing).

The term aligned refers to the starting bit address of a data item relative to certain address boundaries. A starting bit address is block aligned when the low order 7 bits of starting bit address are equal to zero, that is 20 the starting bit address falls on a boundary between adjacent blocks. A word align starting bit address means that the low order 5 bits of starting bit address are zero, the starting bit address points to a boundary between adjacent words. A byte aligned starting bit address means that the low order 3 bits of starting bit address are zero, the starting bit address points to a boundary between adjacent bytes.

u t -Vr r r i r i i j j c -164- Bit granular data has a starting bit address falling within a byte, but not on a byte boundary, or the address is aligned on a byte boundary but the length of the data is bit granular, that is not a multiple of 8 bits.

b. MEM-O--0-2-Phvs-ica-Stt-rt--Pi- 21 Referring to Fig. 201, a partial block diagram of MEM 10112 is shown. Major functional units of MEM 10112 are Main Store Bank (MSB) 20110; including Memory Arrays S(MA's) 20112, Bank Controller (BC) 20114, Memory Cache 10 (MC) 20116, including Bypass Write File (BYF) 20118, Field Isolation Unit (FIU) 20120, and Memory Interface Controller (MIC) 20122.

MSB 20110 comprises MEM 10112's first or bulk level of storage. MSB 20110 may include from one to, for example, 16 MA 20112's. Each MA 20112 may have a storage capacity, for example, 256 K-byte, 512 K-byte, 1 M-byte, or 2 M-bytes of storage capacity. As will be described SC further below, MA 20112's of different capacities may be used together in MSB 20110. Each MA 20112 has a data input connected in parallel to Write Data (WD) Bus 20124 and a data output connected in parallel to Read Data (RD) Bus 20126. MA's 20112 also have control and address ports connected in parallel to address and control (ADCTL) Bus 20128. In particular, Data Inputs 20124 of Memory Arrays 20112 are connected in parallel to Write Data (WD) Bus 20126, and Data Outputs 20128 of Memory Arrays 20112 are connected in parallel to Read Data (RD) Bus 20130.

S''*i -165- Control Address Ports 20132 of Memory Arrays 20112 are connected in parallel to Address and Control (ADCTL) Bus 20134.

Data Output 20136 of Bank Controller 20114 is connected to WD Bus 20126 and Data Input 20138 of BC 20114 is connected to RD Bus 20130. Control and Address Port 20140 of BC 20114 is connected to ADCTL Bus 20134.

BC 20114's Data Input 20142 is connected to MC 20116's Data Output 20144 through Store Back Data (SBD) Bus 10 20146. BC 20114's Store Back Address Input 20148 is connected to MC 20116 Store Back Address Output 20150 through Store Back Address (SBA) Bus 20152. BC 20114's Read Data Output 20154 is connected to MC 20116's Read Data Input 20156 through Read Data Out (RDO) Bus 20158.

BC 20114's Control Port 20160 is connected to Memory Control (MCNTL) Bus 20164.

MC 20116 has Output 20166 connected to MIO Bus 10131 through MIO Port 10128, and Port 20168 connected to MOD Bus 10144 through MJP Port 10140. Control Port 20170 of 20 MC 20116 is connected to MCNTL Bus 20164. Input 20172 of BYF 20118 is connected to IOM Bus 10130 through MIO Port 10128, and Output 20176 is connected to SBD Bus 20146 through Bypass Write In (BWI) Bus 20178.

Finally, FIU 20120 has an Output 20180 and an Input 20182 connected to, respectively, MIO Bus 10129 and IOM Bus 10130 through MIO Port 10128. Input 20184 and Port 20186 are connected to, respectively, JPD Bus 10142 and f MOD Bus 10144 through MJP Port 10140. Control Port 20188 is connected to MCNTL Bus 20164. Referring finally to MIC 20122, MIC 20122 has Control Port 20190 and Input k, 4 -166- 20192 connected to, respectively, IOMC Bus 10131 and IOM Bus 10130 through MIO Port 10128. Control Port 20194 and Input 20196 are connected, respectively, to JPMC Bus 10147 and Physical Descriptor (PD) Bus 10146 through MJP Port 10140. Control Port 20198 is connected-to MCNTL Bus 20164.

c. MEM--O-2--Gener--OpQeptn Referring first to MEM 10112's interface to IOS 10116, this interface includes MIO Bus 10129, IOM Bus 10130, and IOMC Bus 10131. Read and Write Addresses and data to be written into MEM 10112 are transferred from IOS 10116 to MEM 10112 through IOM Bus 10130. Data read from MEM 10112.is transferred to IOS 10116 through MIO Bus 10129. IOMC 10131 is a Bi-directional Control bus between MEM 10112 and IOS 10116 and, as described further below, transfers control signals between MEM 10112 and IOS 10116 to control transfer of data between MEM 10112 t and IOS 10116.

MEM 10112's interface to JP 10114 is MJP Port 10140 20 and includes JPD Bus 10142, MOD Bus 10144, PD Bus 10146, and JPMC Bus 10147. Physical descriptors, that is MEM S10112 physical read and write addresses, are transferred from JP 10114 to MEM 10112 through PD Bus 10146. S Ops, S* that is sequences of S Instructions and operand names, are transferred from MEM 10112 to JP 10114 through MOD Bus 10144 while data to be written into MEM 10112 from JP 10114 is transferred from JP 10114 to MEM 10112 through JPD Bus 10142. JPMC Bus 10147 is a By-directional pl -167- 4* .4 4 4 4 9 4 4 *4 4( if I if i 44 4 I '4 4i* t4 4 Control bus for transferring command and control signals between MEM 10112 and JP 10114 for controlling transfer of data between MEM 10112 and JP 10114. As will be described further below, MJP Port 10140, and in particular MOD Bus 10144 and PD Bus 10146, is generally physically organized as a single port that operates as a dual port. In a first case, MJP Port 10140 operates as a Job Processor Instruction (JI) Port for.transferring S Ops from MEM 10112 to JP 10114. In a second case, MOD 10 10144 and PD 10146 operate as a Job Processor Operand (JO) Port for transfer of operands, from MEM 10112 to JP 10114, while JPD Bus 10142 and PD Bus transfer operands from JP 10114 to MEM 10112.

Referring to MSB 20110, MSB 20110 contains MEM 10112's first, or bulk, level of storage capacity. MSB 20110 may contain from one to, for example, 16 MA's 20112. Each MA 20112 contains a dynamic, random access memory array and may have a storage capacity of, for example 256 Kilo-bytes, 512 Kilo-bytes, 1 Mega-bytes, or 2 Mega-bytes. MEM 10112 may therefore have a physical capacity of up to, for example, 16 Mega-bytes of bulk storage. As will be described further below. MA 20112's of different capacity may be Used together in MSB 20110, for example, four 2 Mega-byte MA 20112's and four 1 Megabyte MA 20112's.

4fl .1 if" -168- BC 20114 controls operation of MA's 20112 and is the path for transfer of data to and from MA's 20112. In addition, BC 20114 performs error detection and correction on data transferred into and out of MA's 20112, refreshes data stored in MA's 20112, and, during a refresh operations, performs error detection and correction of data stored in MA's 20112.

MC 20116 comprises MEM 10112's second, or cache, St. level of storage capacity and contains, for example 8 1, 0 Kilo-bytes of high speed memory. MC 20116, including BYF 20118, is also the path for data transfer between MSB S' 20110 (through BC 20114) and JP 10114 and IOS 10116. In general, all read and write operations between JP 10114 and IOS 10116 are through MC 20116. IOS 10116 may, however, perform read and write operations of complete blocks by-passing MC 20116. Block write operations from IIOS 10116 are accomplished through BYF 20118 while block c read operations are performed through a data transfer path internal to MC 20116 and shown and described below.

S 20 All read and write operations between MEM 10112 and JP 10114, however, must be performed through the cache internal to MC 20116, as will be shown and described further below.

As also shown and described below, FIU 20120 includes write data registers for receiving data to be written into MEM 10112 from JP 10114 and IOS 10116, and circuitry for manipulating data read from MSB 20110 so that MEM 10112 appears as a bit addressable memory. FIU 20120, in

L

-169addition to providing bit addressability of MEM 10112, performs right and left alignment of data, zero fill of data, sign extension operations, and other data manipulation operations described further below. In performing these data manipulation operations on data read from MEM 10112 to JP 10114, MOD Bus 10144 is used as a data path internal to MEM 10112 for transferring of data from MC 20116 to FIU 20120,and from FIU 20120 to MC 20116. That is, data to be transferred to JP 10114 is S. 10 read from MC 20116, transferred through MOD Bus 10144 to FIU 20120, manipulated by FIU 20120, and transferred from FIU 20120 to JP 10114 through MOD Bus 10144.

r t MIC 20122 contains circuitry controlling operation of MEM 10112 and, in particular, controls MEM 10112's interface with JP 10114 and IOS 10116. MIC 20122 L receives MEM 10112 read and write request, that is read Sand write addresses through PD Bus 10146 and IOM Bus 10130 and control signals through JPMC Bus 10147 and IOMC Bus 10131, and provides control signals to BC 20114, MC 20116, and FIU 20120 through MCNTL Bus 20164.

Having described the overall structure and operation S" of MEM 10112, the structure and operation of MEM 10112's Port, MIO Port 10128, and MJP Port 10140, will be described next, followed by descriptions of MEM 10112's control structure and the control and flow of MEM 10112 read and write requests.

I

g i -170d. MEM- -M101--2-- Po rt--Sftrctture MEM 10112 port structure is designed to provide a simple interface to JP 10114 and IOS 10116. While providing fast and flexible operation in servicing MEM 10112 read and write requests from JP 10114 and IOS 10116. In this regard, MEM 10112, as will be described further below, may handle up to 4 read and write requests concurrently and up to, for example, a 63.6 M-byte per second data rate. In addition MEM 10112 is capable of 10 performing bit granular addressing, block read and write operations, and data manipulations, such as alignment and filling, to enable JP 10114 and IOS 10116 to operate most Sefficiently.

*MEM 10112 effectively services requests from three ports. These ports are MIO Port 10128 to IOS 10116, hereafter referred to as IO Port, and JI and JO Ports, described above, to JP 10114. These three ports share the entire address base of MEM 10112, but IOS 10116, for f e example, may be limited from making full use of MEM ?0 10112's address space. Each port has a different set of allowed operations. For example, JO Port can use a bit granular addresses but can reference only 32 bits of data on each request. JI Port can make read requests only to word align 32 bit data items. 10 Port may reference bit granular data, and, as described further below, may read or write up to 16 bytes on each read or write request. i The characteristics of each of these ports will be discussed next below.

I

-171- 1. -O-Pofrt- -Characcte-r-istcs IOS 10116 may access MEM 10112 in either of two modes. The first mode is block transfers by-passing or through the cache in MC 20116, and the second is nonblock transfer through the cache and MC 20116.

Block by-passes may occur for both read and write operations. A read or write operation is eligible for a block by-pass if the data is on block boundaries, is 16 bytes long, and the read or write request is not 10 accompanied by a control signal indicating that an encache (load into MC 20116's cache) operation is to be performed. A by-pass operation takes place only if the block address, that is the physical address of the block in MEM 10112 does not address a currently encached block, that is the block is not present in MC 20116's cache. If the block is encached in MC 20116's cache, the read or write transfer is to MC 20116's cache.

SPartial block references, that is non-full block transfers will go through MC 20116's cache. If a cache miss occurs, that is the reference data is not present in MC 20116's cache, MEM 10112's control structures transfer the data to or from MSB 20110 and update MC 20116's cache. It should be noted that partial blocks may be as short as one byte, or up to 15 bytes long. A starting byte address may be anywhere within a block, but the partial block's length may not cross a block boundary.

I A -172- Bit length transfers, that is transfers of data items having a length of 1 to 16 bits and not a multiple of a byte, or where address is not on a byte boundary, go through MC 20116's cache. These operations may cross byte, word, or block boundaries but may not cross page boundaries. These specific operations requested by port determines whether a read or write request is a partial block or bit length transfer.

2. iTJ-OPort--Chaacte-r-is-tcs S* 10 All read or write requests from JO Port must go So through MC 20116's cache; by-pass operations may not be S performed. The data transferred between MEM 10112 and JP 10114 is always 32 bits in length but, of the 32 bits passed, from zero to 32 bits may be valid data. JP 10114 determines the location of valid data within the 32 bits by referring to certain FIU specification bits provided as part of the read or write request. As will be I described further below, FIU specification bits, and other control bits, are provided to MIC 20122 by JP 10114 20 through JPMC Bus 10147 when each read or write request is made.

While MEM 10112 does not perform block by-pass operations to JP 10114, MEM 10112 may perform a cache 1 read-through operation. Such operations occur on a JP 10114 read request wherein the requested data is not present in MC 20116's cache. If the JP 10114 read request is for a full word, which is word aligned, MEM -173- 10112's Load Manager, discussed below, transfers the requested data directly to JP 10114 while concurrently loading the requested data into MC 20116's cache. This operation is referred to as a "hand-off" operation.

These operations may also be performed by IO Port for 16 bit half words aligned on the right hand half word of a 32 bit word, or if a full block is handed left and loaded into MC 20116's cache.

3. J-1- Port- -eharacterast+cs 10 All JI Port requests are satisfied through MC 20116's Scache; MEM 10112 does not perform by-pass operations to SJI Port. JI Port requests are always read requests for full-word aligned words and are handed off, as described above, if a cache miss occurs: In most other respects, JI Port requests are similar to JO Port requests.

Having described the overall structure and operation of MEM 10112, including MEM 10112's input and output t ports to JP 10114 and IOS 10116, MEM 10112's control structure will be described next below.

J 20 e. MEM" 0-l2G-C+on'1o±- te- p Referring to Fig. 207, a more detailed block diagram of MIC 20116 is shown. Fig. 207 will be referred to in t conjunction with Fig. 201 in the following discussion of MEM 10112's control structure.

1. MEM f-10 l2 G n i -l p e Referring first to Fig. 207, MCNTL Bus 20164 is represented as including MCNTL-BC Bus 20164A, MCNTL-MC Bus 20164B, and MCNTL-FIU Bus 20164C. BUses 20164A, i -174- 20164B, and 20164C are branches of MCNTL Bus 20164 connected to, respectively, BC 20114, MC 20116, and FIU 20120. Also represented in Fig. 207 are PD Bus 10146 and JPMC Bus 10147 to JP 10114, and IOM Bus 10130 and IOMC Bus 10131 to IOS 10116.

JO Port Address Register (JOPAR) 20710 and JI Port Address Register (JIPAR) 20712 have inputs connected from PD Bus 10146. IO Port Address Register'(IOPAR) 20714 has an input connected from IOM Bus 10130. Port Control 0 fl i10 Logic (PC) 20716 has a bi-directional input/outputs connected from JPC 10147 and IOMC Bus 10131. By-pass Read/Write Control Logic (BR/WC) 20718 has a bidirectional input/output connected from IOMC Bus 10131.

Outputs of JOPAR 20710, JIPAR 20712, and IOPAR 20714 are connected to inputs of Port Request Multiplexer (PRMUX) 20720 through, respectively, Buses 20732, 20734, J 20736. PRMUX 20720's output in turn is connected to Bus 20738. Branches of Bus 20738 are connected to inputs of Load Pointers (LP) 20724, Miss Control (MISSC) 20726, and Request Manager (RM) 20722, and to Buses MCNTL-MC 20164B and MCNTL-FIU 20164C.

tSp. ,Outputs of PC 20716 are connected to inputs of JOPAR 20710, JIPAR 20712, IOPAR 20714, PRMUX 20720, and LP 20724 through Bus 20738. Bus 20740 is connected between an input/output of PC 20716 and in input/output of RM 20722. i 1 1 t 1 1 -175- An output of BR/WC 20718 is connected to MCNTL-MC Bus 20164B through Bus 20742. Inputs of BR/WC 20718 are connected from outputs of RM 20722 and Read Queue (RQ) 20728 through, respectively, Buses 20744 and 20746.

RM 20722 has outputs connected to MCNTL-BC Bus 20164A, MCNTL-FIU Bus 20164C, and input of MISSC 20726, and an input of LP 20724 through, respectively, Buses 20748, 20750, 20752, and 20754. MISSC 20726's output is connected to MCNTL-BC Bus 20164A. Outputs of LP 20724 are connected to MCNTL-MC Bus 20164B and to an input of S 10 LM 20730 through, respectively, Buses 20756 and 20758.

0, RQ 20728's input is connected from MCNTL-MC Bus 20164B through Bus 20760 and RQ 20728 has outputs connected to an input of LP 20724, through Bus 20762, and as previously described to an input of BR/WC 20718 through Bus 20746. Finally, LM 20730's output is connected to MCNTL-MC Bus 20164B through Bus 20764.

Having described the structure of MIC 20716 with reference to Fig. 207, and having previously described S' the structure of MEM 10112 with reference to Fig. 201, MEM 10112's control structure operation will next be described with reference to both figures 201 and 207..

t 2. EM-IQ-12--Gono-1-O-peTo-n Referring first to Fig. 207, JOPAR 20710, JIPAR 20712, and IOPAR 20714 are, as previously described, connected from PD Bus 10146 from JP 10114 and 10M Bus 10130 from IOS 10116. JPAR 20710, JIPAR 20712, and IOPAR 20714 receive read and write request addresses from JP 10114 and IOS 10116 and store these addresses for -ii ;L r ii:: -176r I L rt

*F

uA ta I t subsequent service by MEM 10112. As will be described further below, these address inputs from JP 10114 and IOS 10116 include FIU information specifying what data manipulation operations must be performed by FIU 20120 before requested data is transferred to the requestor or written into MEM 10112, information regarding the destination data read from MEM 10112 is to be provided to, information regarding the type of operation to be performed by MEM 10112, and information regarding operand length. Request address ihformation received and stored in JOPAR 20710, JIPAR 20712, and IOPAR 20714 is retained therein until MEM 10112 has initiated service of the corresponding requests. MEM 10112 will accept further request address information into a given port register only after a previous request into that port has been serviced or aborted. Address information outputs from JOPAR 20710, JIPAR 20712, and IOPAR 20714 is transferred through PRMUX 20720 to Bus 20738 and from there to RM 20722, MC 20116, and FIU 20120 as service of individual requests is initiated. As will be described below, this address information will be transferred through PRMUX 20720 and Bus 20738 to LP 20724 for use in servicing a cache miss upon occurrence of a MC 20116 miss.

PC 20716 receives command and control signals pertinent to each requested memory operation from JP 10114 and IOS 10116 through JPMC Bus 10147 and IOSC Bus 10131. PC 20716 includes request arbitration logic and 1 i:;

I

-177port state logic. Request arbitration logic determines the sequence in which IO, JI, JO ports are serviced, and when each port is to be serviced. In determining the sequence of port service, request arbitration logic uses present port state information for each port from the port state logic, information from JPMC Bus 10147 and IOMC Bus 10131 regarding each incoming request, and information from RM 20722 concerning the present state of operation of MEM 10112. Port state logic selects each 10 particular port to be serviced and, by control signals ie through Bus 20738, enables transfer of each port's request address information from JOPAR 20710, JIPAR 20712, and IOPAR 20714 through PRMUX 20720 to Bus 20738 •*Ic for use by the remainder of MEM 10112's control logic in servicing the selected port. In addition to request information received from JP 10114 and IOS 10116 through *JPMC Bus 10147 and IOMC Bus 10131, port state logic utilizes information from RM 20722 and, upon occurrence of a cache miss, from LM 20730 (for clarity of 20 presentation, this connection is not represented in Fig.

207). Port state logic also controls various port state flag signals, for example port availability signals, signals indicating valid requests, and signals indicating *ffffC that various ports are waiting service.

RM 20722 controls execution of service for each request. RM 20722 is a microcode controlled "micromachine" executing programs called for by requested MEM 10112 operations. Inputs of RM 20722 include request address information from IOPAR 20714, JIPAR 20212, and .3 -178- JOPAR 20210, including information regarding the type of MEM 10112 operation to be performed in servicing a particular request, interrupt signals from other MEM 10112 control elements, and, for example, start signals from PC 20716's request arbitration logic. RM 20722 provides control signals to FIU 20120, MC 20116, and most other parts of MEM 10112's control structure.

Referring to Fig. 201, MC 20116's cache is, for example, an 8 Kilo-byte, four set associative cache used to provide rapid access to a subset of data stored in MSB 20110. The subset of MSB 20110 data stored in MC 20116's bache at any time is the data most recently used by JP 10114 or IOS 10116. MC 20116's cache, described further below, includes tag store comparison logic for determining encached addresses, a data store containing corresponding encached data, and registers and logic necessary to up-date cache contents upon occurrence of a cache miss. Registers and logic for servicing cache misses includes logic for determining the least recently CV t used cache entry and registers for capture and storage of S 20 information regarding missed cache references, for t example modify bits and replacement page numbers. Inputs "o to MC 20116 are provided from RM 20722, LM 20730 (discussed further below), FIU 20120, MSB 20110 (through BC 20114), LP 20724 (described further below) and address information from PRMUX 20720. Outputs of MC 20116 include data and go to FIU 20120 (through MOD Bus 10144), -179the data requestors (JP 10114 and IOS 10116), and a MC 20116 Write Back File (described further below).

As previously described, FIU 20120 includes logic necessary to make MEM 10112 appear bit addressable. In addition, FIU 20120 includes logic for performing certain data manipulation operations as required by the requestors (JP 10114 or IOS 10116). Data is transferred into FIU 20120 from MC 20116 through that portion of MOD Bus 10144 internal to MEM 10112, is manipulated as required, and is then transferred to the requestor through MOD Bus 10144 or MIO Bus 10129. In the case of writes requiring read-modify-write of encached data, the data is transferred back to MC 20116 through MOD Bus 10144 after manipulation. In general, data manipulation operations include locating requested data onto selected MOD Bus 10144 or MIO Bus 10139 lines and filling unused r bus lines as specified by the requestor. Data inputs to FIU 20120 may be provided from MC 20116 or JP 10114 through MOD Bus 10144 or from IOS 10116 through IOM Bus 20 10130. Data outputs from FIU 20120 may be provided to MC 20116, JP 10114, or IOS 10116 through these same buses.

Control information is provided to FIU 20120 from RM S. 20722 through Bus 20748 and MCNTL-FIU Bus 20164C.

Address information may be provided to FIU 20120 from JOPAR 20710, JIPAR 20712, or IOPAR 20714 through PRMUX 20720, Bus 20738, and MCNTL-FIU Bus 20164C.

t I2 S. -180- Returning to Fig. 207, MISSC 20726 is used in handling MC 20116 misses. In the event of a request referring to data not in MC 20116's cache, MISSC 20726 stores block address of the reference and type of operation to be performed, this information being provided from an address register in MC 20116 and from RM 20722. MISSC 20726 utilizes this information in generating a command to BC 20114, through MCNTL-BC Bus 20164A, for a data read from MSB 20110 to obtain the referenced data. BC 20114 places this command in a queue, or register, and subsequently executes the commanded read operation. MISSC 20726 also generates an entry into RQ 20728 (described further below) indicating l the type of operation to be performed when referenced data is subsequently read from MSB 20110.

RQ 20728 is, for example, a three-level deep queue Sstoring information indicating operations associated with .O.t data being read from MSB 20110. Two kinds of operation may be indicated: block by-pass reads and cache loads.

20 If a cache load is specified, that is a read nd store to MC 20116' cache, .is indicated, RM 20722 is interrupted and forced to place other MEM 10112 operations in idle S. until cache load is completed. A block by-pass read operation results in by-pass read control (described below) assuming control of the data from MSB 20110.

C Inputs to RQ 20728 are control signals from RM 20752, MISSC 20726, and BC 20114. RQ 20728 provides control 4' *1 Pr -181- 4 *r 9 ii 4 4c outputs to LP 20724 (described below) LM 20730 (described below) RM 20722, and by-pass read control (described below).

LP 20724 is a set of registers for storing information necessary for servicing MC 20116 misses that result in order to load MC 20116's tag store. LM 20730 uses this information when data stored in MSB 20110 and read from MSB 20110 to service a MC 20116 cache miss, becomes available through BC 20114. Inputs to LP 20724 include the address of the missing reference, provided from JOPAR 20710, JIPAR 20712, or IOPAR 20714 through PRMUX 20720 and Bus 20738, commands from RM 20722, and a control signal from RQ 20728. LP 20724 outputs include addresses of missed references to MC 20116, through Bus 15 20756 and MNCTL-MC 20164B, and command signals to LM 20730 and BR/WC 20718.

LM 20730, referred to above, controls loading of MC 20116's cache with data from MSB 20110 after occurrence of a cache miss. RQ 20728, referred to above, indicates, 20 for each data read from MSB 20110, whether the data read is the result of a MC 20116 cache miss. If the data is read from MSB 20110 as a result of a cache miss, LM 20730 proceeds to issue a sequence of control signals for loading the data from MSB 20110 and its associated address into MC 20116's cache. This data is transferred into MC 20116's cache data store while the block address, from LP 20724 is transferred into the tag store

I.

4 44..

9.

*c I 19 94 tc I 4 4..

i -182- (described in the following discussion) of MC 20116's cache. If the transfer of data into MC 20116's cache replaces data previously resident in that cache, and that previous data is "dirty", that is has been written into so as to be different from an original copy of the data stored on MSB 20110, the modified data resident in MC 20116's cache must be written back into MSB 20110. This operation is performed through a Write Back File contained in MC 20116 and described below. In the event of such an operation, LM 20730 initiates a write back operation by MC 20116 and BC 20114, also as described below.

As will be described further in a following description, all MC 20116 cache load operations are full a 4 word blocks. A request resulting in a MC 20116 cache miss may result in a "hand-off", that is a read operation of a full 4 word block. Handoff operations also may be of single 32 bit words wherein a 32 bit word aligned word is transferred from JP 10114 or a 16 bit operand aligned on the right half-word is transferred from IOS 10116. In 20 such a handoff operation, LM 20730 will send a valid request signal to the requesting port and a handoff S. operation will be performed. Otherwise, a waiting signal will be sent to the requesting port and the request will re-enter the priority queue of PC 20716 for subsequent execution. To accomplish these operations, LM 20730 receives input from RQ 20728, (not shown in E'ig. 207 for S4| r -183- 'clarity of presentation) and LP 20724. LM 20730 provides outputs to port state logic of PC 20716, to MC 20116, MC 20116's Write Back File and MC 20116's Write Back Address Register and to BC 20114.

Referring to Fig. 201, as previously discussed IOS 20116 may request a full block write operation directly to MSB 20110. Such a by-pass write request may be honored if the block being transferred is not encached in MC 20116's cache. In such a case, RM 20722 will initiate the transfer setting up By-Pass Write Control logic in BR/WC 20718, and may then pass control of the operation over to BR/WC 20718's By-Pass Write Control logic for completion. By-Pass Write Control may then accept the remaining portion of the data block from IOS 10116, j 15 generating appropriate hand shaking signals through IOMC Bus 10131, and load the data block into BYF 20118 and MC 20116. MISSC 20726 will provide a by-pass write command to BC 20114, through MNCTL-PC Bus 20164A. BC 20114 will then tansfer the data block from BYF 20118 and into MA's 20 20112 and MSB 20110.

As previously described, BYF 20118 receives data from IOM Bus 10130 and provides data output to BC 20114 1 through BWY Bus 20178 and SBD Bus 20146. BYF 20118 is capable of simultaneously accepting data from IOM Bus 25 10130 while reading data out to BC 20114. Control of l c writing data into YF 20118 is provided from BR/WC go,. 20718's By-Pass Write Control logic.

L. 1 1 t 1 i I- j -184- IOS 10116 may, as previously described, request a full block read operation by-passing MC 20116's cache.

In such a case, BR/WC 20718's by-pass read control handles data transfer to IOS 10116 and generates required hand shaking signals to IOS 10116 through IOMC Bus 10131. The data path for by-pass read operations is through a data path internal to MC 20116, rather than through BYF 20118. This internal data path is RDO Bus 20158 to MIO Bus 10129.

As previously described, BC 20114 manages all data transfers to and from MA's 20112 in MSB 20110. BC 20114 receives requests for data transfers from RM 20722 in an internal queue register. All data transfers to and from MSB 20110 are full block transfers with block aligned 15 addresses. On data write operations, BC 20114 receives data from BWF 20118 or from MC 20116's Write Back File and transfers the data into MA's 20112. During read t operations, BC 20114 fetches the data block from MA's 20112 and places the data block on RDO Bus 20158 while 20 signalling to MIC 20122 that the data is available. As S, described above, MIC 20122 tracks and controls transfer of data and BYF 20118, MC 20116, and MC 20116's Write Back File, and directs data read from MSB 20110 to the appropriate destination, MC 20116's Data Store, JP 10114, or IOS 10116.

In addition to the above operations, BC 20114 ,controls refresh of MA's 20112 and performs error detection and correction operations. In this regard, BC 20114 performs two error detection and correction 1 1 -185operations. In the first, BC 20114 detects single and double bit errors in data read from MSB 20110 and corrects single bit errors. In the second, BC 20114 reads data stored in MA's 20112 during refresh operations and performs single bit error detection. Whenever an error is detected, during either read operations or refresh operations, BC 20114 makes a record of that error in an error log contained in BC 20114 (described further in a following description). Both JP 10114 and IOS 10116 may read BC 20114's error log, and information from BC 20114's error log may be recorded in a CS 10110 maintenance log and to assist in repair and trouble shooting of CS 10110. BC 20114's error log may be caddressed directly by RM 20722 and data from BC 20114's error log is transferred to JP 10114 or IOS 10116 in the same manner as data stored in MSB 20110.

SReferring finally to MA's 20112, each MA 20112

S

t contains an array of dynamic semiconductor random access memories. Each MA 20112 may contain 256 Kilo-bytes, 512 Kilo-bytes, 1 Mega-bytes, or 2 Mega-bytes of data :.storage. The storage capacity of each MA 20112 is organized as segments of 256 Kilo-bytes each. In addressing a particular MA 20112, BC 20114 selects that particular MA 20112 as will be described further below.

BC 20114 concurrently selects a segment within that MA ,r 20112, and a block of four words within that segment.

Each word may comprise 39 bits of information, 32 bits of data and 7 bits of error correcting code. The full 39 bits of each MA 20112 word are transferred between BC x- 1 -186- 20114 and MA's 20112 during each read and write operation. Having briefly described the general structure and operation of MEM 10112, certain types of operations which may be performed by MEM 10112 will be described next below.

f. MEM--1OQ1-2-Operat!io MEM 10112 may perform two general types of operation. The first type are data transfer operations and the second type are memory maintenance operations.

Data transfer operations may include read, write, and read and set. Memory maintenance operations may include read error log, repair block, and flush cache. Except during a flush cache operation, the existence of MC 20116 and its operation is invisible to the requestors, that is JP 10114 and IOS 10116.

A MEM 10112 read operation transfers data from MS S10112 to a requestor, either JP 10114 or IOS 10116. A 4 read data transfer is asynchronous in that the requestor cannot predict elapsed time between submission of a memory operation request and return of requested data.

Operation of a requestor in MEM 10112 is coordinated by a S" requested data available signal transmitted from MEM .W 10112 to the requestor.

4 |A MEM 10112 write operation transfers data from either JP 10114 or IOS 10116 to MEM 10112. During such S, operations, JP 10114 is not required to wait for a signal from MEM 10112 that data provided to MEM 10112 from JP i Et .4

L

-187- 10114 has been accepted. JP 10114 may transfer data to MEM 10112's JO Port whenever a JO Port available signal from MEM 10112 is present; read data is accepted immediately without further action or waiting required of JP 10114. Word write operations from IOS 10116 are performed in a similar manner. On block write operations, however, IOS 10116 is required to wait for a data taken signal from MEM 10112 before sending the 2nd, 3rd and 4th words of a block.

MEM 10112 has a capability to perform "lock bit" operations. In such operations, a bit granular read of the data is performed and the entire operand is t transmitted to the requestor. At the same time, the most significant bit of the operand, that is the Lock Bit, is set to one in the copy of data stored in MEM 10112. In the operand sent to the requestor, the lock bit remains at its previous value, the value before the current read and set operation. Test and set operations are performed by performing read and set operations wherein the data 20 item length is specified to be one bit.

4* As previously described, MEM 10112 performs certain S..maintenance operations, including error detection, MEM 10112's Error Log in BC 20114 is a 32 bit register *containing an address field and an error code field. On a first error to occur, the error type and in some cases, such as ERCC errors on read data stored in MSB 20110, the address of the data containing the aeror are stored in BC 20114's Error Log Register. An interrupt signal -188indicating detection of an error is raised at the same that information regarding the error is stored in the Error Log. If multiple errors occur before Error Log is read and reset, the information regarding the first error will be retained and will remain valid. The Error Log code field will, however, indicate that more than one error has occurred.

JP 10114 may request a read Error Log operation referred to as a "Read Log and Reset" operation. In this operation, MEM 10112 reads the entire contents of Error Log to JP 10114, resets Error Log Register, and resets the interrupt signal indicating presence of an error.

*IOS 10116, as discussed further below, is limited to reading 16 bits at a time from MEM 10112. It therefore S.4 15 requires two read operations to read Error Log. First read operation to IOS 10116 reads an upper 16 bits of Error Log data and does not reset Error Log. The second read operation is performed in the same manner as a JP 10114 Read Log and Reset operation, except that only the 20 low order 16 bits of Error Log are read to IOS 10116.

MEM 10112 performs repair block operations to correct parity or ERCC errors in data stored in MC 20116's Cache or in data stored in MA's 20112. In a repair block procedure, parity bits for data stored in MC 20116's Cache, or ERCC check bits of data stored in MA's 20112, are modified to agree with the data bits of data stored therein. In this regard, repaired uncorrectible errors, such as two bit errors of data in MA's 20112, will have good ERCC and parity values. Until a repair block -189operation is performed, any read request directed to bad data, that is data having parity or ERCC check bits indicating invalid data, will be flagged as invalid.

Repair block operations therefore allow such data to be read as valid, for example to be used in a data correction operation. Errors are ignored and not logged in BC 20114's Error Log in repair block operations. A write operation into an area containing bad data may be accomplishe, if MEM 10112's internal operation does not require a read-modified-write procedure. Only byte aligned writes of integral byte length data residing in MC 20116 and word aligned writes of integral word lengths of data in MSP 20110 do not require read-modified-write operation. By utilizing such write operations, it is therefore possible to overwrite bad data by use of normal write operations before or instead of repair block operations.

"f MEM 10112 performs a cache flush operation in event of a power failure, that is when MEM 10112 goes into battery back-up operation. In such an event, only MA's S 20112 and BC 20114 remain powered. Before JP 10114 and *OS 10116 lose power, JP 10114 and IOS 10116 must transfer to MEM 10112 any data, including operating state, to be saved. This is accomplished by using a series of normal write operations. After conclusion of these write operations, both JP 10114 and IOS 10116 transmit a flush cache request to MEM 10112. Upon

IJ,

L i 1 1 1 1 1 -190receiving two flush cache requests, MEM 10112 flushes MC 20116's Cache so that all dirty data encached in MC 20116's Cache is transferred into MA's 20112 before power is lost. If only JP 10114 or IOS 10116 is operating, DP 10118 will detect this fact and will have transmitted an enabling signal (FLUSHOK) to MEM 10112 during system initialization. FLUSHOK enables MEM 10112 to perform cache flush upon receiving a single flush cache request.

After a cache flush operation, no further MEM 10112 operations are possible until DP 10118 resets a power failure lock-out signal to enable MEM 10112 to resume normal operation.

"i I Having described MEM 10112's overall structure and #1 operation and certain operations which may be performedby MEM 10112, MEM 10112's interfaces to JP 10114 and IOS 10116 will be described next below.

t 01 -2-Inte~faces-to-J-+0l a4-id-IeS 01 o6-^P-gs--2 .2120--2---2 As previously described, MJP Port 10140 and MIO Port 10128 logically function as three independent ports.

These ports are an I1 Port to IOS 10116, a JP Operand Port to JP 10114 and a JP Instruction Port to JP 10114.

Referring to Figs. 209, 210, and 211, diagramic representations of IO Port 20910, JP Operand (JPO) Port 21010, and JP Instruction (JPI) Port 21110 are shown respectively.

Port 20910 handles all IOS 10116 requests to MEM 10112, including transfer of both instructions and i -191operands. JPO Port 21010 is used for read and write operations of operands, for example numeric values, to and from JP 10114. JPI Port 21110 is used to read SINs, that is SOPs and operand NAMEs, from MEM 10112 to JP 10114. Memory service requests to a particular port are serviced in the order that the requests are provided to the Port. Serial order is not maintained between requests to different ports, but ports may be serviced in the order of their priority. In one embodiment of the 1, 0 present invention, 10 Port 20910 is accorded highest priority, followed by JPO Port 21010, and lastly by JPI Port 21110, with requests currently contained in a port having priority over incoming requests. As described above and will be described in more detail in following descriptions, MEM 10112 operations are pipelined. This pipelining allows interleaving of requests from 10 Port 20910, JPO Port 21010, and JPI Port 21110, as well as overlapping service of requests at a particular port. By overlapping operations it is meant that one operation servicing a particular port begins before a previous operation servicing that port has been completed.

I Cha aet JsTt^i- f

J

FCt^- -2 fr9-^ SReferring first to Fig. 209, a diagramic representation of 10 Port 20910 is shown. Signals are transmitted between 10 Port 20910 and IOS 10116 through MIO Bus 10129, IOM Bus 10130, and IOMC Bus 10131. MIO 4 -192- Bus 10129 is a unidirectional bus having inputs from MC 20116 and FIU 20120 and dedicated to transfers of data and instructions from MEM 10112 to IOS 10118. IOM Bus 10130 is likewise a unidirectional bus and is dedicated to the transfer, from IOS 10118 to MEM 10112, of read addresses, write addresses, and data to be written into MEM 10112. IOM Bus 10130 provides inputs to BYF 20118, FIU 20120, and MIC 20122. IOMC Bus 10131 is a set of dedicated signal lines for the exchange of control 10 signals between IOS 10118 and MEM 10112.

Referring first to MIO Bus 10129, MIO Bus 10129 is a 36 bit bus receiving read data inputs from MC 20116's Cache and from FIU 20120. A single read operation from SMEM 10112 to IOS 10116 transfers one 32 bit word (or 4 bytes) of data (MIO(0-31)) and four bits of odd parity S(MIOP(0-3)), or one parity bit per byte.

Referring next to IOM Bus 10130, a single transfer from IOS 10116 to MEM 10112 includes 36 bits of S" information which may comprise either a memory request comprising a physical address, a true length, and command bits. These memory requests and data are multiplexed onto IOM 10130 by IOS 10116.

Data transfers from IOS 10116 to MEM 10112 each comprise a single 32 bit data word (IOM(0-31)) and four bits of odd parity (IOMP(0-3)) or one parity bit per byte. Such data transfers are received by either BYF 20118 or FIU 20120.

-193- Each IOS 10116 memory request to MEM 10112, as described above, an address field, a length field, and an operation code field. Address and length fields occupy the 32 IOM Bus 10130 lines used for transfer of data to MEM 10112 in IOS 10116 write operations. Length field includes four bits of information occupying bits (IOM(0of IOM Bus 10130 and address field contains 27 bits of information occupying bits (IOM(4-31)) of IOM Bus v 10130. Together, address and length field specify a 10 physical starting address and true length of the c particular data item to be written into or read from MEM 10112. Operation code field specifies the type of operation to be performed by MEM 10112. Certain basic operation codes comprise 3 bits of information occupying bits (IOMP (32-36)) of IOM Bus 10130; as described above. These same lines are used for transfer of parity bits during data transfers. Certain operations which may be requested of MEM 10112 by IOS 10116 are, together with 'l their corresponding command code fields, are; 000 read, 001 read and set, 010 write, c 011 error, 100 read error log (first half), 101 read error log (second half) and reset, 110 repair block, and 111 flush cache.

-194- Two further command bits may specify further operations to be performed by MEM 10112. A first command bit, indicates to MEM 10112 during write operations whether it is desirable to encache the data being written into MEM 10112 in MC 20116's Cache. IOS 10116 may set this bit to zero if reuse of the data is unlikely, thereby indicating to MEM 10112 that MEM 10112 should avoid enchaching the data. IOS 10116 may set this bit to one if the data is likely to be reused, thereby 10 indicating to MEM 10112 that it is preferable to encache the data. A second command bit is referred to a CYCLE.

CYCLE command bit indicates to MEM 10112 whether a t particular data transfer is a single cycle operation, that is a bit granular word, or a four cycle operation, that is a block aligned block or a byte aligned partial t" block.

IOMC 10131 includes a set of dedicated lines for exchange of control signals between IOS 10116 and MEM 10112 to coordinate operation of IOS 10116 and MEM 10112. A first such signal is Load IO Request (LIOR) from IOS 10116 to MEM 10112. When IOS 10116 wishes to load a memory request into MEM 10112, IOS 10116 asserts LIOR to MEM 10112. IOS 10116 must assert LIOR during the same system cycle during which the memory request, that is address, length, and command code fields, are valid.

If LIOR and 10 Port Available (IOPA) signals, described below, are asserted during the same clock cyle, MEM 10112's port is loaded from IOS 10116 and IOPA is -195dropped, indicating the request has been accepted. If a load of a request is attempted and IOPA is not asserted, MEM 10112 remains unaware of the request, LIOR remains active, and the request must then be repeated when IOPA is asserted.

IOPA is a signal from MEM 10112 to IOS 10116 which is asserted by MEM 10112 when MEM 10112 is available to accept a new request from IOS 10116. IOPA may be asserted while a previous request from IOS 10116 is completing operation if the address, length, and S, operation code fields of the previous request are no longer required by MEM 10112, for example in servicing bypass operations.

IO Data Taken (TIOMD) is a signal from MEM 10112 to IOS 10116 indicating that MEM 10112 has accepted data from IOS 10116. IOS 10116 places a first data word on IOM Bus 10130 on the next system clock cycle after a write request is loaded; that is, LIOR has been asserted, a memory request presented, and IOPA dropped. MEM 10112 then takes that data word on the clock edge beginning the next system clock cycle. At this point, MEM 10112 asserts TIOMD to indicate the data has been accepted. On it a single word operations TIOMD is not used by IOS 10116 as a first data word is always accepted by MEM 10112 if 10 Port 20910 was available. On block operations, a first data word is always taken but a delay may occur between acceptance of first and second words. IOS 10116 I 1 L -196is required to hold the second word valid on IOM Bus 10130 until MEM 10112 responds with TIOMD to indicate that the block operation may proceed.

Data Available for IO (DAVIO) is a signal asserted by MEM 10112 to IOS 10116 indicating that data requested by IOS 10116 is available. DAVIO is asserted by MEM 10112 during the system clock cycle in which MEM 10112 places the requested data on MIO Bus 10129. In any single word *type transfer, DAVIO is active for a single system clock S 10 transfer. In block type transfers, DAVIO is normally i active for four consecutive system clock cycles. Upon event of a single cycle "bubble" resulting from detection and correction of an ERCC error by BC 20114, DAVIO will e remain high for four non-consecutive system clock cycles and with a single cycle bubble, a non-assertion, in DAVIO corresponding to the detection and correction of the error.

IO Memory Interrupt (IMINT) is a signal asserted by .MEM 10112 to IOS 10116 when BC 20114 places a record of a detected error in BC 20114's Error Log, as described above.

Previous MIO Transfer Invalid (PMIOI) signal is similarly a signal asserted by MEM 10112 to IOS 10116 regarding errors in data read from MEM 10112 to IOS 10116. If an uncorrectible error appears in such data, that is an error in two or more data bits, the incorrect data is read to IOS 10116 anl PMIOI signal asserted by -197- MEM 10112. Correctible, or single bit, errors in data do not result in assertion of PMIOI. MEM 10112 will assert PMIOI to IOS 10116 of the next system clock cycle following MEM 10112's assertion of DAVIO.

Having described MEM 10112's interface to IOS 10116, and certain operations which IOS 10116 may request of MEM 10112, certain MEM 10112 operations within the capability of the interface will be described next. First, operand transfers, for example of numeric data, between MEM 10112 and IOS 10116 may be bit granular with any length from 10 one to sixteen bits. Operand transfers may cross n boundaries within a page but may not cross physical page 1 boundaries. As previously described, MIO Bus 10129 and IOM Bus 10130 are capable of transferring 32 bits of data at a time. The !east significant 16 bits of these buses, i 15 that is bits 16 to 31, will contain right justified data during operand transfers. The contents of the most significant 16 bits of these buses is generally not defined as MEM 10112 generally does not perform fill oerations on read operations to 10 Port 20910, nor does IOS 10116 fill unused bits during write operations.

During a read or write operation, only those data bits indicated by length field in the corresponding memory request are of significance. In all cases, however, parity must be valid on all 32 bits of MIO Bus 10129 and IOM Bus 10130.

i -198- Referring to Fig. 204, IOS 10116 includes Data Channels 20410 and 20412 each of which will be described further in a following detailed description of IOS 10116. Data Channels 20410 and 20412 each possess particular characteristics defining certain 10 Port 20910 operations. Data Channel 20410 operates to read and write block aligned full and partial blocks. Full blocks have block aligned addresses and lengths of 16 bytes.

V Partial blocks have byte aligned addresses and lengths of 1 to 15 bytes; a partial block transfer must be within a block, that is not cross block boundaries. A full 4 word block will be transferred between IOS 10116 and MEM 10112 in either case, but only those blocks indicated by length of field in a corresponding MEM 10112 request are of actual significance in a write operation. Non-addressed bytes in such operations may contain any information so cr- long as parity is valid for the entire data transfer.

Data Channel 20412 preferably reads or writes 16 bits at a time on double byte boundaries. Such reads and writes are right justified on MIO Bus 10129 and IOM Bus 10130.

The most significant 16 bits of these buses may contain any information during such operations so long as parity is valid for the entire 32 bits. Data Channel 20412 operations are similar to IOS 10116 operand read and write operations with double byte aligned addresses and lengths of 16 bits. Finally, instructions, for example controlling IOS 10116 operation, are read from MEM 10112 -199to IOS 10116 a block at a time. Such operations are identical to a full block data read.

Having described the operating characteristics of Port 20910, the operating characteristics of JPO Port 21010 will be described next.

2. -JtPO--Pot-2 O-O--peratin j hatacteristics-+Piir--2IO Referring to Fig. 210, a diagramic representation of JPO Port 21010 is shown. As previously described, JPO Port 21010 is utilized for transfer of operands, for example numeric data, between MEM 10112 and JP 10114.

JPO Port 21010 includes a request input (address, length, and operation information) to MIC 20122 from 36 bit PD Bus 10146, a write data input to FIU 20120 from 32 bit JPD Bus 10142, a 32 bit read data output from MC 20116 and FIU 20120 to 32 bit MOD Bus 10144, and bi-directional control inputs and outputs between MIC 20122 and JPMC Bus 10147.

Referring first to JPO Port 21010's read data output to MOD Bus 10144, MOD Bus 10144 is used by JPO Port 21010 to transfer data, for example operands, to JP 10114. MOD Bus 10144 is also utilized internal to MEM 10112 as a biv, ;directional bus to transfer data between MC 20116 and FIU 20120. In this manner, data may be transferred from MC 20116 to FIU 20120 where certain data format operations are performed on the data before the data is transferred to JP 10114 through MOD Bus 10144. Data may also be used to transfer data from FIU 20120 to MC 20116 after a data format operation is performed in a write operation, Data L L. -200- **99 9 .9i 0 9 *i 0 *Ci may also be transferred directly from MC 20116 to JP 10114 through MOD Bus 10144. Internal to MEM 10112, MOD Bus 10144 is a 36 bit bus for concurrent transfer of 32 bits of data, MOD Bus 10144 bits (MOD(0-31)), and 4 bits of odd parity, 1 bit per byte, MOD Bus 10144 bits (MODP(0-3)). External to MEM 10112, MOD Bus 10144 is a 32 bit bus, comprising bits (MOD(0-31)); parity bits are not read to JP 10114.

Data is written into MEM 10112 through JPD Bus 10142 10 to FIU 20120. As just described, data format operations may then be performed on this data before it is transferred from FIU 20120 to MC 20116 through MOD Bus 10144. In such operations, JPD Bus 10142 operates as a 32 bit bus carrying 32 bits of data, bits (JPD with no parity bits. JO Port 21010 generates parity for JPD Bus 10142 data to be written into MEM 10112 as this data is transferred into MEM 10112.

Memory requests are also transmitted to MEM 10112 from JP 10114 through JPD Bus 10142, which operates in this regard as a 40 bit bus. Each such request includes an address field, a length field, an FIU field specifying data formating operations to be performed, operation code field, and a destination code field specifying destination of data read from MEM 10112. Address field includes a 13 bit physical page number field, (JPPN(0and a 14 bit physical page offset field, (JPPO(0- Length field includes 6 bits of length

I

2 99I 0 4 944499 «rt I -201information, (JLNG(0-5)), and expresses true length of the data item to be written to or read from MEM 10112.

As JPD Bus 10142 and MOD Bus 10144 are each capable of transferring 32 bits of data in a single MEM 10112 read or write cycle, 6 bits of length information are required to express true length. As will be described in a following description, JP 10114 may provide physical page offset and length information directly to MEM 10112, performs logical page number to physical page'number S 10 translations, and may perform a Protection Mechanism 10230 check on the resulting physical page number. As o.r such, MEM 10112 expects to receive (JPPN(0-12)) later than (JPPO(0-13)) and (JLNG(0-5)). (JPPO(0-13)) and should, however, be valid during the system clock cycle in which a JP 10114 memory request is loaded A t into MEM 10112.

'I t Operation code field provided to MEM 10112 from JP 10114 is a 3 bit code, (JMCMD(0-2)) specifying an operation to be formed by MEM 10112. Certain operations which JP 10114 may request of MEM 10112, and their corresponding operato±in codes, are: 000 read; 4 ttt 001 read and set; 010 write; 011 error; 100 error; 101 read error log and reset; 110 repair block; and, 111 flush cache.

-202- Two bit FIU field, (JFIU(0-1)) specifies data manipulation operations to be performed in executing read and write operations. Among the data manipulation operations which may be requested by JP 10114, and their FIU fields, are: 00 right justified, zero fill; 01 right justified, sign extend; 10 left justify, zero fill; and, 11 left justify, blank fill.

For write operations, JPO Port 21010 may respond only to the most significant bit of FIU field, that is the FIU field bit specifying alignment.

Finally, destination field is a two bit field specifying a JP 10114 destination for data read from MEM 10112. This field is ignored for write operations to MEM S10112. A first bit of destination field, JPMDST, identifies the destination to be FU 10120, and the second field, EBMDST, specifies EU 10120 as the destination.

SJPMC Bus 10147 includes dedicated lines for exchange of control signals between JPO Port 21010 and JP 10114.

Among these control signals is Load JO Request (LJOR), which is asserted by JP 10114 when JP 10114 wishes to 'load a request into MEM 10112. LJOR is asserted concurrently with presentation of the memory requ .st to MEM 10112 through PD Bus 10146. JO Port Available (JOPA) is asserted by MEM 10112 when JPO Port 21010 is available to accept a new memory request from JP 10114. If LJOR and JOPA are asserted concurrently, MEM 10112 accepts the memory request from JP 10114 and MEM 10112 drops JOPA to indicate that memory request has been accepted. As r II I I ~p~ -203previously discussed, MEM 10112 may assert JOPA while a previous request is being executed and the PD Bus 10146 information, that is the memory request previously provided concerning the previous request, is no longer required.

If JP 10114 submits a memory request and JOPA is not asserted by MEM 10112, MEM 10112 does not accept the request and JP 10114 must resubmit that request when JOPA is asserted. Because, as described above, JPPN field of a memory request from JP 10114 may arrive late compared to the other fields of the request, MEM 10112 will delay 4loading of JPPN field for a particular request until the next system clock cycle after the request was initially submitted. MEM 10112 may also obtain this JPPN field at the same time it is being loaded into the port register by by-passing the port register.

JP 10114 may abort a memory request upon asserting Abort JP Request (ABJR). ABJR will be accepted by MEM 10112 during system clock cycle after accepting memory 20 request from JP 10114 and ABJR will result in cancellation of the requested operation. A single ABJR line is provided for both JPO Port 21010 and JPI Port 21110 because, as described in a following description, MEM 10112 may accept only a single request from JP 10114, to either JPO Port 21010 or to JPI Port 21110, during a single system clock cycle.ii r -204o 9 9 o 9 O* 9 9 9999 99t 99 9 99.9 9 9 99 9 6 ii *o Upon completion of an operand read operation requested through JPO Port 21010 MEM 10112 may assert either of two data available signals to JP 10114. These signals are data available for FA(DAVFA) and data available for EB(DAVEB). As previously described, a part of each read request from JP 10114 includes a destination field specifying the intended destination of the requested data. As will be described further below, MEM 10112 tracks such destination information for read 10 requests and returns destination information with a corresponding information in the form of DAVFA and DAVEB. DAVFA indicates a destination in FU 10120 while DAVEB indicates a destination in EU 10122. MEM 10112 may also assert signal zero filled (ZFILL) specifying whether read data for JPO Port 21010 is zero filled. ZFILL is valid only when DAVEB is asserted.

For JPO Port 21010 write request, the associated write data word should be valid on same system clock cycle as the request, or one system clock cycle later.

20 JP 10114 asserts Load JP Write Data (LJWD) during the system clock cycle when JP 10114 places valid write data on JPD Bus 10142.

As previously discussed, when MEM 10112 detects an error in servicing a JP 10114 request MEM 10112 places a record of this error in MC 20116's Error Log. When an entry is placed in Error Log for either JPO Port 21010 or IO Port 20910, MEM 10112 asserts an interrupt flag signal Si v -205- *004 *0 *s 0 0* *0 Str t it 0- 0 00 0 ii 0 1 0 I 1£ indicating a valid Error Log entry is present. DP 10118 detects this flag signal and may direct the flag signal to either JP 10114 or IOS 10116, or both. IOS 10116 or JP 10114, as selected by DP 10118, may then read and reset Error Log and reset the flag. The interrupt flag signal is not necessarily directed to the requestor, JP 10114 or IOS 10116, whose request resulted in the error.

If an uncorrectible MEM 10112 errorr that is an error in two or more bits of a single data word, is detected in a read operation the incorrect data is read to JP 10114 and an invalid data signal asserted. A signal, Previous MOD Transfer Invalid (PMODI), is asserted by MEM 10112 on the next system clock cycle following either DAVFA or DAVEB. PMODI is not asserted for single bit errors, 15 instead the data is corrected and the corrected data read to JP 10114.

Having described JPO Port 21010's structure, and characteristics, JPI Port 21110 will be described next below.

20 3. JP--Pot--2+-Operatin Characteisties-fPit--W^ Referring to Fig. 211, a diagramic representation of JPI Port 21110 is shown. JPI Port 21110 includes an address input from PD Bus 10146 to FIU 20120, a data output to MOD Bus 10144 from MC 20116, and bi-directional control inputs and outputs from MIC 20122 to JPMC Bus 10147. As previously described, a primary function of JPI Port 21110 is the transfer of SOPs and operand NAMEs from MEM 10112 to JP 10114 upon request from JP 10114.

£0to 0 1 IF I r( r -206- 4 4 4 e 4

I

*s .5 It SI £4 4 II JPI Port thereby performs only read operations wherein each read operation is a transfer of a single 32 bit word having a word aligned address.

Referring to JPI Port 21110 input from PD Bus 10146, read requests to MEM 10112 by JP 10114 for SOPs and operand NAMEs each comprise a 21- bit word address. As described above, each JPI Port 21110 read operation is of a single 32 bit word. As such, the five least significant bits of address are ignored by MEM 10112.

10 For the same reason, a JPI Port 21110 request to MEM 10112 does not include a length field, an operation code field, an FIU field, or a destination code field.

Length, operation code, and FIU code fields are not required since JPI Port 21110 performs only a single type 15 of operation and destination code field is not required because destination is inherent in a JPI Port 21110 request.

The 32 bit words read from MEM 10112 in response to JPI Port 21110 requests are transferred to JP 10114 through MC 20116's 32 bit output to MOD Bus 10144. As in the case of JPO 21010 read outputs to JP 10114, JPI Port 21110 does not provide parity information to JP 10114.

Control signals exchange between JP 10114 and JPI Port 21110 through JPMC Bus 10147 include Load JI Request 25 (LJIR) and JI Port Available (JIPA), which operate in the same manner as discussed with reference to JPO Port 21010. As previously described, JPO Port 21010 and JPI

I

4 1 -207- Port 21110 share a single Abort JP Request (ABJR) command. Similarly, JPO Port 21010 and JPI Port 21110 share Previous MOD Transfer Invalid (PMODI) from MEM 10112. As described above, a JPI Port 21110 request does not include a destination field as destination is implied. MEM 10112 does, however, provide a Data Available Signal (DAVFI) to JP 10114 when a word read *11.

from MEM 10112 in response to a JPI Port 21110 request is present on MOD Bus 10144 and valid.

10 Having described the overall structure and operation of MEM 10112, and the structure and operation of MEM 10112's interface to JP 10114 and IOS 10116, the structure and operation of FIU 20120 MEM 10112 will next be described in further detail.

15 h. F'IUO-2O--20-F-icgs-;-2-.---2--2) l As previously described, FIU 20120 performs certain Sdata manipulation operations, including those operations necessary to make MEM 10112 bit addressable. Data manipulation operations may be performed on data being 20 written into MEM 10112, for example, JP 10114 through JPD Bus 10142 or from IOS 10116 through IOM Bus 10130. Data Smanipulations operations may also be performed on data being read from MEM 10112 to JPD 10114 or IOS 10116. In t, case of data read to JP 10114, MOD Bus 10144 is used both as a MEM 10112 internal bus, in transferring data from MC 20116 to FIU 20120 for manipulation, and to transfer 1 1 -208manipulated data from MEM 10112 to JP 10114. In case of data read to IOS 10116, MOD Bus 10144 is again used as MEM 10112 internal bus to read data from MC 20116 to FIU 20120 for subsequent manipulation. The manipulated data is then read from FIU 20120 to IOS 10116 through MIO Bus 10129.

:f .Certain data manipulation operations which may be performed by FIU 20120 have been previously described.

In general, a data manipulation operation consists of 10 four distinct operations, and FIU 20120 may manipulate data in any possible manner which may be achieved through performing any combination of these operations. These four possible operations are selection of data to be manipulated, rotation or shifting of that data, masking of that data, and transfer of that manipulated data to a selected destination. Each FIU 20120 data input will comprise a thirty-two bit data word and, as described above, may be selected from input provided from JPD Bus 10142, MOD Bus 10144, and IOM Bus 10130. In certain t+ 20 cases, an FIU 20120 data input may comprise two thirtytwo bit words, for example, when a cross word operation is performed generating an output comprised of bits from each of two different thirty-two bit words. Rotation or i \shifting of a selected thirty-two bit data word enables bits within a selected word to be repositioned with respect to word boundaries. When used in conjunction with the masking operation, described momentarily, rotation and shifting may be reiterably performed to fl -209- C 4 ft i *a 4 kRa3 ?t a ~J a a a 4 a o transfer any selected bits in a word to any selected locations in that word. As will be described further below, a masking operation allows any selected bits of a word to be affectively erased, thus leaving only certain other selected bits, or certain selected bits to be forced to predetermined values. A masking operation may be performed, for example, to zero fill or sign extend portions of a thirty-two bit word. In conjunction with a rotation or shifting operation, a masking operation may, 10 for example, select a single bit of a thirty-two bit input word, position that bit in any selected bit location, and force all other bits of that word to zero.

Each output of FIU 20120 is a thirty-two bit data word and, as described above, may be transferred on to MOD Bus 15 10144 or onto MIO Bus 10129. As will be described below, selection of a particular sequence of the above four operations to be performed on a particular data word is determined by control inputs provided from MIC 20122.

These control inputs from MIC 20122 are decoded and executed by microinstruction control logic included within FIU 20120.

Referring to Fig. 230, a partial block diagram of FIU 20120 is shown. As indicated therein, FIU 20120 includes Data Manipulation Circuitry (DMC) 23010 and FIU Control Circuitry (FIUC) 23012. Data Manipulation Circuitry 23010 in turn includes FIUIO circuitry (FIUIO) 23014, Data Shifter (DS) 23016, Mask Logic (MSK) 23018, and Assembly Register (AR) logic 23020. Data manipulation ci

I

-210- 99 4,L 9.

4,( r99 n.* 4 4,I 4, 4 4,4 4, 9 9.

4, 4 circuitry 23010 will be described first followed by FIUC 23012. In descrbing data manipulation circuitry 23010, FIUIO 23014 will be described first, followed by DS 23016, MSK 23018, and AR 23020, in that order.

Referring to FIUIO 23014, FIUIO 23014 comprises FIU 20120's data input and output circuitry. Job Processor Write Data Register (JWDR) 23022, IO System Write Data Register (IWDR) 23024, and Write Input Data Register (RIDR) 23026 are connected from, respectively, JPD Bus 10 10142, IOM Bus 10130, and MOD Bus 10144 for receiving data word inputs from, respectively, JP 10114, IOS 10116, and MC 20116. *JWDR 23022, IWDR 23024 and RIDR 23026 are each thirty-six bit registers comprised, for example, of SN74S374 registers. Data words transferred into IWDR 15 23024 and RIDR 23026 are each, as previously described, comprised of a thirty-two data word plus four bits of parity. Data inputs from JP 10114 are, however, as previously described, thirty-two bit data words without parity. Job Processor Parity Generator (JPPG) 23028 associated with JWDR 23022 is. connected from JPD Bus 10142 and generates four bits of parity for each data input to JWDR 23022. JWDR 23022's thirty-six bit input thereby comprises thirty-two bits of data, directly from JPD Bus 10142, plus a corresponding four bits of parity from JPPG 23028.

4,494 9 4, 4 4,CI *944,,4 WWW I- -211- Data words, thirty-two bits of data plus four bits of parity, are transferred into JWDR 23022, IWDR 23024, or RIDR 23026 when, respectively, input enable signals Load JWD (LJWD), Load IWD (LIWD) or Load RID (LRID) are asserted. LJWD is provided from FU 10120 while LIWD and LRID are provided from MIC 20122.

Data words resident in JWDR 23022, IWDR 23024, or RIDR 23026 may be selected and transferred onto FIU 20120's Internal Data (IB) Bus 23030 by output enable 10 signals JWD Enable Output (JWDEO), IWD Enable Output (IWDEO), an RID Enable Output (RIDEO). JWDEO, IWDEO, and RDIEO are provided from FIUC 23012 described below.

S As will be described further below, manipulated data words from DS 23016 or AR 23020 will be transferred onto, respectively, Data Shifter Output (DSO) Bus 23032 or Assembly Register Output (ASYRO) Bus 23034 for subsequent transfer onto MOD Bus 10144 or MIO Bus 10129. Each manipulated data word appearing on DSO Bus 23032 or ASYRO Bus 23034 will be comprised of 32 bits of data plus 4 bits of parity. Manipulated data words present on DSO Bus 23032 may be transferred onto MOD Bus 10144 or MIO Bus 10129 through, respectively, DSO Bus To MOD Bus Driver Gate (DSMOD) 23036 or BSO Bus To MIO Bus Driver Gate (DSMIO) 23038. Manipulated data words present on ASYRO Bus 23034 may be transferred onto MOr Bus 10144 or MIO Bus 10129 through, respectively, ASYRO Bus To MOD Bus Driver Gate (ASYMOD) 23040 or ASYRO Bus To MIO Bus Driver Gate (ASYMIO) 23042. DSMOD 23036, DSMIQ 23038, ASYMOD ;j -212- 23040, and ASYMIO 23042 are each comprised of, for example, SN74S244 drivers. A manipulated data word on DSO Bus 23032 be transferred through DSMOD 23036 to MOD Bus 10144 when driver gate enable signal Driver Shift To MOD (DRVSHFMOD) to DSMOD 23036 is asserted. Similarly, a manipulated data word on DSO Bus 23032 will be transferred through DSMIO 23038 to MIO Bus 10129 when driver gate enable signal Drive Shift Through MIO Bus (DRVSHFMIO) to DSMIO 23038 is asserted. Manipulated data words present on ASYRO Bus 23034 may be transferred onto MOD Bus 10144 or MIO Bus 10129 when, respectively, driver gate enable signal Drive Assembly To Mod Bus (DRVASYMOD) *4 to ASYMOD 23040 or Drive Assembly To MIO Bus (DRVASYMIO) to ASYMIO 23042 are asserted. DRVSHFMOD, DRVSHFMIO, DRVASYMOD, and DRVASYMIO are provided, as described Sbelow, from FIUC 23012.

Registers IARM 23044 and BARMR 23046, which will be la described further in a following description of DP 10118, are used by DP 10118 to, respectively, write data words onto IB 23030 and to Read data words from MOD Bus 10144, St for example manipulated data words from FIU 20120. Data ,I Sord written into IARMR 23044 from DP 10118, that is 32 bits of data and 4 bits of parity, will be transferred onto IB Bus 23030 when register enable output signal IARM enable output (IARMEO) from FIUC 23012 is asserted.

Similarly, a data word present on MOD Bus 10144, comprising 32 bits of data plus 4 bits of parity, will be written into BARR 23046 when load enable signal Load -213- BARMR (LDBARMR) to BARMR 23046 is asserted by MIC 20122.

A data word written into BARMR 23046 from MOD Bus 10144 may then subsequently be read to DP 10118. IARMR 23044 and BARMR 23046 are similar to JWDR 23022, IWDR 23024, and IRDR 23026 and may be comprised, for example, of SN74S299 registers.

Referring finally to IO Parity Check Circuit (IOPC) 23048, IOPC 23048 is connected from IB Bus 23030 to receive each data word, that is 32 bits of data plus 4 0- 10 bits of parity, appearing on IB Bus 23030. IOPC 23048 confirms parity and data validity of each data word appearing on IB Bus 23030 and, in particular, determines validity of parity and data of data words written into FIU 20120 from IOS 10116. IOPC 23048 generates output 15 Parity Error (PER), previously discussed, indicating a i. parity error in data words from IOS 10116.

Referring to DS 23016, DS 23016 includes Byte Nibble Logic (BYNL) 23050, Parity Rotation Logic (PRL) 23052, and Bit Scale Logic (BSL) 23054. BYNL 23050, PRL 23052, and BSL 23054 may respectively be comprised of, for example, 25S10 shifters. BYNL 23050 is connected from IB Bus 23030 for receiving and shifting the 32 data bits of a data word selected and transferred onto IB Bus 23030.

PRL 23052 is a 4 bit register similarly connected from IB Bus 23030 to receive and shift the 4 parity bits of a data word selected and transferred onto IB Bus 23030.

Outputs of BYNL 23050 and PRL 23052 are both connected onto DSO Bus 23032, thus providing a 36 bit FIU 20120 -214data word output directly from BYNL 23050 and PRL 23052.

BYNL 23050's 32 bit data output is also connected to BSL 23054's input. BSL 23054's 32 bit output is in turn provided to MSK 23018.

As previously described, DS 23016 prforms data manipulation operations involving shifting of bits within a data word. In general, data shift operations performed by DS 23016 are rotations wherein data bits are right shifted, with least significant bits of data word being shifted into most significant bit position and most significant bits being translated towards least significant bit positions. DS 23016 rotation operations are performed in two stages. First stage is performed by BYNL 23050 and PRL 23052 and comprises right rotations on S. 15 a nibble basis (a nibble is defined as 4 bits of data).

That is, BYNL 23050 right shifts a data word by an integral number of 4 bit increments. A right rotation on a nibble by nibble basis may, for example, be performed when RM 20722 asserts FLIPHALF previously described.

FLIPHALF is asserted for IOS 10116 half word read -operations wherein the request data resides in the most significant 16 bits of a data wor, from MC 20116. 3YNL 23050 will perform a right rotation of 4 nibbles to transfer the desired 16 bits of data into the least significant 16 bits of BYNL 23050's output. Resulting BYNR 23050 output, together PRL 23052's parity bit output would then be transferred through DSO 23050 to MIO Us 10129. In addition to performing data shifting

-I

-215operations, DS 23016 may transfer a data word, that is the 32 bits of data, directly to MSK 23018 when data manipulation to be performed does not require data shifting, that is shifts of 0 bits may be performed.

Because data bits are shifted by BYNL 23050 on a nibble basis, the relationship between the 32 data bits of a word and the corresponding 4 parity bits may be a. maintained if parity bits are similarly right rotated by an amount corresponding to right rotation of data bits.

This relationship is true if the data word is shifted in multiples of 2 nibbles, that is 8 bits or 1 byte. PRL 23052 right rotates the 4 parity bits of a data word by I an amount corresponding to right rotation of the corresponding 32 data bits in BYNL 23050. Right rotated outputs of BYNL 23056 and PRL 23052 therefore comprise a S valid data word having 32 bits of data and 4 bits of parity wherein the parity bits are correctly related to the't data bits. A right rotated data word output from BYNL 23050 and PRL 23052 may be transferred onto DSO Bus 23032 for subsequent transfer to MOD Bus 10144 or MIO Bus i 10129 as described above. DSO 23032 is used as FIU S20120's output data path for byte write operations and S* "rotate read" operations wherein the required manipulation of a particular data word requires only an integral numer of right rotations by bytes. Amount of right rotation of 32 bits of data in BYNL 23050 and 4 bits of parity in PRL 23052 is controlled by input signal shift (SHFT) to BYNL 23050 and PRL 23052. As will -216be described below, SHFT is generated, together with SHFT controlling BSL 23054, by FIUC 23012.

BYNL 23050 and PRL 23052, like BSL 23054 described below, are parallel shift logic chips and entire rotation operation of BYNL 23050 and PRL 23052 or BSL 23054 may be performed in a single clock cycle.

Second stage of rotation is performed by BSL 23054 which, as described above, receives the 32 data bits of a data word from BYNL 23050. BSL 23054 performs right 0 rotation on a bit by bit basis with the shift amount being selectable between 0-3 bits. Therefore, BSL 23054 ^a may rotate bits through nibble boundaries. BYNL 23050 and BSL 23054 therefore comprise a data shifting circuit capable of performing bit-by-bit right rotation by an 15 amount from 1 bit to a full 32 bit right rotation.

Referring now to MSK 23018, MSK 23018 is comprised of 32 bit Mask Word Generators (MWG's) 23056 to 23064.

MSK 23018 generates a 32 bit output to AR 23020 by selectively combining 32 bit mask word outputs of MWG's 23056 to 23064. Each mask word generated by one of MWG's i 23056 to 23064 is effectively comprised a bit by bit combination of a set of enabling bits and a pre- S" determined 32 bit mask word, generated by FIUC 23012 and MIC 20122. MWG's 23058 to 23064 are each comprised of for example, open collector NAND gates for performing these functions, while NWG 23056 is comprised of a PROM.

-u:ls l~ l C- -217- As just described, outputs of MWG's 23056 to 23064 are all open collector circuits so that any selected combination of mask word outputs from MWG's 23056 to 23064 may be ORed together to comprise the 32 bit output of MSK 23018.

MWG 23056 to MWG 23064 generate, respectively, mask word outputs Locked Bit Word (LBW) Sign Extended Word (SEW) Data Mask Word (DMW) Blank Fill Word (BWF) and Assembly Register Output 10 (ARO) Referring first to MWG 23064 and ARO (0- 31), the contents of Assembly Register (ASYMR) 23066 in AR 23020 are passed through MWG 23064 upon assertion of enabling signal Assembly Output Register (ASYMOR) ARO (0-31) is thereby a copy of the contents of ASYMR 23066 and MWG 23064 allows the contents of ASYMR 23066 to be f C ORed with the selected combination of LBW SEW (0- 31), DMW or BFW (0-31).

DMW (0-31) from MWG 23060 is generated by ANDing enable Input Data Mask (DMSK) (0-31) with the 32 bit output of DS 23016. DMSK (0-31) is a 32 bit enabling word generated, as described below, by FIUC 23012. FIUC 23012 may generate 4 different DMSK (0-31) patterns.

Referring to Fig. 231, the 4 DMSKs (0-31) which may be generated by FIUC 20132 are shown. DMSKA (0-31) is shown in Line A of Fig. 231. In DMSKA (0-31) all bits to the left of but not including a bit designated by Left Bit Address (LBA) and all bits to the right of and not including a bit designated by Right Bit Address (RBA) are c i -218- 0. All bits between, and including, those bits designated by LBA and RBA are l's. DMSKB (0-31) is shown in Line B of Fig. 231 and is DMSKA (0-31) inverted.

DMSKC (0-31) and DMSKD (0-31) are shown, respectively, in Lines C and D of Fig. 231 and are comprised of, respectively, all 0's or all l's. As stated above DMSK (0-31) is ANDed with the 32 bit output of DS 23016. As such, DMSKC (0-31) may be used, for example, to inhibit DS 23016's output while DMSKD (0-31) may be used, for S o. 10 example, to pass DS 23016's output to AR 23020. DMSKA (0-31) and DMSKB (0-31) may be used, for example, to gate selected portions of DS 23016's output to AR 23020 where, r" for example, the selected portions of DS 23016's output may be ORed with other mask word outputs MSK 23018.

Referring next to MWG 23062, MWG 23062 generates BFW t t BFW (0-31) is used in a particular operation wherein 32 bit data words containing 1 to 4 ASCII blanks are required to be generated wherein 1 bit/byte contains a logic one and remaining bits contain logic zeros. In this case, the ASCII blank bytes may contain logic l's in bit positions 2, 10, 18, and 26.

Referring again to Fig. 231, Line E therein shows 32 bit right mask (RMSK) (0-31) which may be generated by FIUC 23012. In the most general case, RMSK contains zeros in all bit positions to the left of and including a bit position designated by RBA. When used in a blank fill operation, bit positions 2, 10, 18, and 26 may be selected to contain logic l's depending upon those byte positions containing logic 1's, that is in those bytes i^ li L ^J -219containing ASCII blanks; these bytes to the right of RBA are determined by RMSK RMSK (0-31) is enabled through MWG 23062 as BWF (0-31) when MWG 23062 is enabled by blank fill (BLNKFILL) provided from FIU 23012.

As described above, MWG's 23058 to 23064 and in particular MWG's 23060 and MWG 23062 are NAND gate operations. Therefore, the outputs of MWGs 23056 through 23064 are active low signals. The inverted output of ASYMR 23066 is used as an output to ASYRO 23034 to invert 10 these outputs to active high.

MWG 23058, generating SEW is used in generating sign extended or filled words. In sign extended words, all bit spaces to the left of the most significant bit of a 32 bit data word are filled with the sign bit of the data contained therein, the left most bits of the 32 bit word are filled with 1's or O's *depending on whether that word's sign bit indicates that the data contained therein is positive or negative number.

Sign Select Multiplexor (SIGNSEL) 23066 is connected to receive the 32 data bits of a word present on IB Bus 23030. Sign Select (SGNSEL) to SIGNSEL 23066 is derived from SBA that is from SBA Bus 21226 from PRMUX 20720. As previously described, SBA is Starting Bit Address identifying the first or most significant bit of a data word. When a data word contains a signed number, most significant bit contains sign bit of that number. SGNSEL input to SIGNSEL 23066 is used as a selection input and, when SIGNSEL is -220enabled by Sign Extend (SIGNEXT) from FIU 23012, selects the sign bit on IB Bus 23030 and provides that sign bit as an input to MWG 23058.

Sign bit input to MWG 23058 is ANDed with each bit of left hand mask (LMSK) (0-31) from FIUC 23012. Referring again to Fig. 231, LMSK (0-31) is shown on Line F thereof. LMSK (0-31) contains all O's to the right of and including the bit space identified by LBA and l's in all bit spaces to the left of that bit space identified by LBA. SEW (0-31) will therefore contain sign bit in all bit spaces to the left of the most significant bit of the data word present on output of MWG 23058. The data Swocd on IB Bus 23030 may then be passed through DS 23016 and subjected to a DMSK operation wherein all bits to the left of the most significant bit are forced to 0. SEW (0-31) and DMW (0-31) outputs of MWG's 23058 and 23060 may then be ORed to provide the desired find extended word output.

LBW provided by MWG 23056, is used in locked bit operations wherein the most significant data bit of a data word is in MEM 10112 forced to logic 1. SIGNSEL (0- 4) is an address input to MWG 23056 and, as previously described, indicates most significant data bit of a data word present on an IB Bus 23030. MWG 23056 is enabled by input Lock (LOCK) from FIUC 23012 and the resulting LBW (0-31) will contain a single logic 1 in the bit space of the most significant data bit of the data word present on IB Bus 23030. The data word present on IB Bus 23030 may

A

-221then be passed through DS 23016 and MWG 23060 to be ORed with LBW (0-31) so that that data words most significant data bit is forced to logic 1.

Referring to AR 23020, AR 23020 includes ASYMR 23066, which may be comprised for example of a SN74S175 registers, and Assembly Register Parity Generator (ASYPG) 23070. As previously described, ASYMR 23066 is connected from MSK 23018 32 bit output. A 32 bit word present on MSK 23018's output will be transferred into ASYMR 23066 when ASYMR 23066 is enabled by Assembly Register Load (ASYMLD) from MIC 20122. The 32 bit word generated through DS 23016 and MSK 23018 will then be present on ASYRO Bus 23034 and may, as described above, then be transferred onto MOD Bus 10144 or MIO Bus 10129. ASYPG 23070 is connected from ASYMR 23066 32 bit output and will generate 4 parity bits for the 32 bit word presently on the 32 data lines of ASYRO Bus 23034. ASYPG 23070's 4 bit parity output is bused on the 4 parity bit lines of S*ASYRO Bus 23034 and accompany the 32 bit data word present thereon.

Having described structure and operation of Data Manipulation Circuitry 23010, FIUC 23012 will be described next below.

Referring again to Fig. 230, FIUC 23012 provides pipelined microinstruction control of FIU 20120. That is, control signals are received from MIC 20122 during a first clock cycle and certain of the control signals are decoded by microinstruction logic to generate further ,i i i iC -221then be passed through DS 23016 and MWG 23060 to be ORed with LBW (0-31) so that that data words most significant data bit is forced to logic 1.

Referring to AR 23020, AR 23020 includes ASYMR 23066, which may be comprised for example of a SN74S175 registers, and Assembly Register Parity Generator (ASYPG) 23070. As previously described, ASYMR 23066 is connected from MSK 23018 32 bit output. A 32 bit word present on MSK 23018's output will be transferred into ASYMR 23066 when ASYMR 23066 is enabled by Assembly Register Load (ASYMLD) from MIC 20122. The 32 bit word generated through DS 23016 and MSK 23018 will then be present on ASYRO Bus 23034 and may, as described above, then be transferred onto MOD Bus 10144 or MIO Bus 10129. ASYPG 23070 is connected from ASYMR 23066 32 bit output and will generate 4 parity bits for the 32 bit word presently on the 32 data lines of ASYRO Bus 23034. ASYPG 23070's 4 bit parity output is bused on the 4 parity bit lines of ASYRO Bus 23034 and accompany the 32 bit data word present thereon.

Having described structure and operation of Data Manipulation Circuitry 23010, FIUC 23012 will be described next below.

Referring again to Fig. 230, FIUC 23012 provides pipelined microinstruction control of FIU 20120. That is, control signals are received from MIC 20122 during a first clock cycle and certain of the control signals are decoded by microinstruction logic to generate furthert i -222- 99 9 9 0 9.

0 9 0*4 4 44 44 4 4 44.4 4 I 1/1 C 4t 94 4CE 4.4*4*9 .4 r 9 FIUC 23012 control signals. During the second clock cycle, control signals received and generated during first clock cycle are provided to DMC 23010, some of which are further decoded to provide yet other control signals to control operation of FIUC 23012. FIUC 23012 includes Initial Decode Logic (IDL) 23074, Pipeline Registers (PPLR) 23072, Final Decoding Logic (FDL) 23076, and Enable Signal Pipeline Register (ESPR) 23098 with Enable Signal Decode Logic (ESDL) 23099.

10 IDL 23074 and Control Pipeline Register (CPR) 23084 of PPLR 23072 are connected from control outputs of MIC 20122 to receive control signals therefrom during a first clock cycle as described above. IDL 23074 provides outputs to control pipeline registers Right Bit Address Register (RBAR) 23086, Left Bit Address Register (LBAR) 23088 and Shift Register (SHFR) 23090 of PPLR 23072. CPR 23084 and SHFR 23090 provide control outputs directly to DMC 23010. As described above these outputs control DMC 23010 during second clock cycle, 20 CPR 23084, RBAR 23086, and LBAR 23088 provide outputs to FDL 23076 during second clock cycle and FDL 23076 in turn provides certain outputs directly to DMC 23010.

ESPR 23098 and ESDL 23099 receive enable and control signals from MIC 20122 and in turn provide enable and control signals to DMC 23010 and certain other portions of MEM 10112 circuitry.

i

$I

-223- IDL 23074 and FDL 23076 may be comprised, for example, of PROMs. CPR 23084, RBAR 23086, LBAR 23088, SHFR 23090, and ESPR 23098 may be comprised, for example, of SN74S194 registers. ESDL 23099 may be comprised of, for example, compatible decoders, such as logic gates.

Referring first to IDL 23074, IDL 23074 performs an initial decoding of circuitry control signals from MIC "20122 and provides further control signals used by FIUC 23012 in controlling FIU 20120. IDL 23074 is comprised 0. 10 of read-only memory arrays Right Bit Address Decoding Logic (RBADL) 23078, Left Bit Address Decoding Logic (LBADL) 23080, and Shift Amount Decoding Logic (SHFAMTDL) 23082. RBADL 23078 receives, as address inputs, Final Bit Address (FBA) Bit Length Number (BLN) and Starting Bit Address (SBA) FBA, BLN and SBA r define, respectively, the final bit, length, and starting V" bit of a requested data item as previously discussed with reference to PRMUX 20720. RBADL 23078 also receives chip select enable signals Address Translation Chip Select (ATCS) 00, 01, 02, 03, 04, and 15 from MIC 20122 and, in particular, RM 20722. When FIU 20120 is required to execute certain MSK 23018 operations, inputs FBA BLN and SBA together with an ATCS input, y are provided to RBADL 23078 from MIC 20122. RBADL 23078 t in turn provides output RBA (Right Bit Address) which has been described above with reference to DMSK (0- 31) and RMSK LBADL 23080 is similar to RBADL 23078 and is provided with inputs BLN FBA -224- SBA and ATCS 06, 07, 08, 09, and 05 from MIC 20122. Again, for certain MSK 23018 operations, LBADL 23080 will generate Left Bit Address (LBA) which has been previously discussed above with reference to DMSK (0-31) and LMSK (0-31).

RBA and LBA are, respectively, transferred to RBAR 23086 and LBAR 23088 at start of second clock cycle by Pipeline Load Enable signal PIPELD provided from MIC 20122. RBAR 23086 and LBAR 23088 in I 10 turn respectively provide outputs Register Right Address .(RRAD) and Register Left Address (RLAD) as address inputs to Right Mask Decode Logic (RMSKDL) 23092, Left Mask Decode Logic (LMSKDL) 23094, and FDL 23076 at start of second clock cycle. RRAD and RLAD (0-4) S 15 correspond respectively to RBA and LBA RMSKDL 23092 and LMSKDL 23094 are ROM arrays, having, as just described, RRAD and RLAD as, respectively, address inputs and Mask Enable (MSKENBL) from CPR 23084 as enable inputs. Together, RMSKDL 23092 and LMSKDL 23094 generate, respectively, RMSK (0-31) and LMSK (0-31) to MSK 23018. RMSK (0-31) and LMSK (0-31) are provided as inputs to Exclusive Or/Exclusive Nor gating (XOR/XNOR) 23096. XOR/XNOR 23096 also-receives enable and selection signal Out Mask (OUTMSK) from CPR 23084. RMSK (0-31) and LMSK (0-31) inputs to XOR/XNOR 23096 are used, as selected by OUTMSK from CPR 23084, to generate a selected DMSK (0-31) as shown in Fig. 231.

DMSK (0-31) output of XOR/XNOR 23096 is provided, as described above, to MSK 23018.

ppr -225sr .4, 04 9 *i 4 *i *4 .4 .4..C *4 4. 4 *i 0 44 4:*y a *l a Referring again to IDL 23074, SHFAMTBL 23082 decodes certain control inputs from MIC 20122 to generate, through SHFR 23090, control inputs SHFT and SGNSEL to, respectively, DS 23016, SIGNSEL 23068 and MWG 23056. Address inputs to the PROMs comprising SHFAMTBL 23082 include FBA SBA and FLIPHALF (FLIPHALF) from MIC 20122. FBA and SBA have been described above. FLIPHALF is a control signal indicating that, as described above, that 16 bits of data requested by IOS 10116 resides in the upper half of a 32 bit data word and causes those 16 bits to be transferred to the lower half of FIU 20120's output data word onto MIO Bus 10129. MIC 20122 also provides chip enable signals ATCS 10, 11, 12, 13, and 14. Upon receiving these control inputs from MIC 20122, SHFAMTDL 23082 generates an output shift amount (SHFAMT) which, together with SBA from MIC 20122, is transferred into SHFR 23090 by PIPELD at start of second clock cycle. SHFR 23090 then provides corresponding outputs SHFT and SIGNSEL As described above, SIGNSEL are provided to SIGNSEL 23068 and MWG 23056 and MSK 23018. SHFT is provided as SHFT and SHFT to, respectively, BYNL 23050 and BSL 23054 and DS 23016.

Referring to CPR 23084, as described above certain control signals are provided directly to FIU 20120 circuitry without being decoded by IDL 23074 or FDL 23076. Inputs to CPR 23084 include Sign Extension >4 -226- (SIGNEXT) and Lock (LOCK) indicating, respectively, that FIU 20120 is to perform a sign extension operation through MWG 23058 or a lock bit word operation through MWG 23056. CPR 23084 provides corresponding outputs SIGNEXT and LOCK to MSK 23018 to select these operations. Input Assembly Output Register (ASYMOR) and Blank Fill (BLANKFILL) are passed through CPR 23084 as ASYMOR and BLANKFILL to, respectively, MWG 23064 and MWG "23062 to select the output of ASYMR 23066 as a mask or to 10 indicate that MSK 23018 is to generate a blank filled word through MWG 23062. Inputs OUTMSK and MSKENBL to CPR S23084 are provided, as discussed above, as enable signals .o OUTMSK and MSKENBL to, respectively, EXOR/ENOR 23096 and RMSKDL 23092 and LMSKBL 23094 and generating RMSK (0-31), LMSK and DMSK (0-31) as described above.

Referring finally to ESPR 23098 and ESDL 23099, ESPR 23098 and PPLR 23072 together comprise a pipeline register and ESDL 23099 decoding logic for providing enable signals to FIU 20120 and other MEM 10112 circuitry. ESPR 23098 receives inputs Drive MOD Bus (DRVMOD) Drive MIO Bus (DRVMIO) and Enable Register (ENREG) from MIC 20122 as previously "r described. DRVMOD DRVMIO and ENREG (0-1) are transferred into ESPR 23098 by PIPELD as previously described with reference to PPLR 23072. ESPR 23098 provides corresponding outputs to ESDL 23099, which in turn decodes DRVMOD DRVMIO and ENREG (0-1) to provide enable signals to FIU 20120 and other MEM -227- 10112 circuitry. Outputs DRVSHFMOD, DRVASYMOD, DRVSHFMIO, and DRVASYMIO are provided to DSMOD 23036, DSMIO 23038, ASYMOD 23040, ASYMIO 23042, and FIUIO 23014 to control transfer of FIU 20120 manipulated data words onto MOD Bus 10144 and MIO Bus 10129. Outputs IARMEO, JWDEO, IWDEO, and RIDEO are provided as output enable signals to IARMR 23044, JWDR 23022, IWDR 23024, and RIDR *l 23026 to transfer the contents of these registers onto IB Bus 23030 as previously described. Outputs DRVCAMOD, 10 DRVAMIO, DRVBYMOD, and DRVBYMIO are provided to MC 20116 for use in controlling transfer of information onto MOD 4, ,Bus 10144 and MIO Bus 10129.

Having described the structure and operation of MEM 10112 above, the structure and operation of FU 10120 will S. 15 be described next below.

B. F etch--Un-it--e--2 f4s:--20-2- 6 10 4 As has been previously described, FU 10120 is an independently operating, microcode controlled machine comprising, together with EU 10122, CS 10110's micromachine for executing user's programs. Principal 9:,o functions of FT 10120 include: Fetching and interpreting instructions, that is SINs comprising SOPs and Names, and data from MEM 10112 for use by FU 10120 and EU 10122; Organizing and controlling flow and execution of user programs; Initiating EU 10122 operations; Performing arithmetic and logic operations on data; Controlling transfer of data from

A-

-228- FU 10120 and EU 10122 to MEM 10112; and, Maintaining certain stack register mechanisms. Among these stack and register mechanisms are Name Cache (NC) 10226, Address Translation Cache (ATC) 10228, Protection Cache (PC) 10234, Architectural Base Registers (ABRs) 10364, Micro- Control Registers (mCRs) 10366, Micro-Stack (MIS) 10368, Monitor Stack (MOS) 10370 of General Register File (GRF) 10354, Micro-Stack Pointer Register Mechanism (MISPR) ¢e 10356, and Return Control Word Stack (RCWS) 10358. In 10 addition to maintaining these FU 10120 resident stack and register mechanisms, FU 10120 generates and maintains, in S' whole or part, certain MEM 10112 resident data S"IC structures. Among these MEM 10112 resident data structures are Memory Hash Table (MHT) 10716 and Memory 15 Frame Table (MFT) 10718, Working Set Matrix (WSM) 10720, t Virtual Memory Management Request Queue (VMMRQ) 10721, Active Object Table (AOT) 10712, Active Subject Table (AST) 10914, and Virtual Processor State Blocks (VPSBs) 10218. In addition, a primary function of FU 10120 is the generation and manipulation of logical descriptors which, as previously described, are the basis of CS 10110's internal addressing structure. As will be described further below, while FU 10120's internal structure and operation allows FU 10120 to execute arithmetic and logic operationq, FU 10120's structure includes certain features to expedite generation and manipulation of logical descriptors.

-229- Referring to Fig. 202, a partial block diagram of FU 10120 is shown. To enhance clarity of presentation, certain interconnections within FU 10120, and between FU 10120 and EU 10122 and MEM 10112 are not shown by line connections but, as described further below, are otherwise indicated, such as by common signal names.

Major functional elements of FU 10120 include Descriptor Processor (DESP) 20210, MEM 10112 Interface Logic Oa(MEMINT) 20212, and Fetch Unit Control Logic (FUCTL) 0: 10 20214. DSP 20210 is, in general, an arithmetic and logic unit for generating and manipulating entries for MEM S 10112 and FU 10120 resident stack mechanisms and caches, as described above, and, in particular, for generation and manipulation of logical descriptors. In addition, as 15 stated above, DSP 20210 is a general purpose Central 9*fe Processor Unit (CPU) capable of performing certain S' arithmetic and logic functions.

DESP 20210 includes AON Processor (AONP) 20216, Offset Processor (OFFP) 20218, Length Processor (LENP) 20220. OFFP 20218 comprises a general, 32 bit CPU with additional structure to optimize generation and manipulation of offset fields of logical descriptors.

AONP 20216 and LENP 20220 comprise, respectively, iprocessors for generation and manipulation of AON and length fields of logical descriptors and may be used in conjuction with OFFP 20218 for execution of certain i arithmetic and logical operations. DESP 20210 includes GRF 10354, which in turn include Global Registers (GRs) 10360 and Stack Registers (SRs) 10362. As previously J i -230described, GR's 10360 includes ABRs 10364 and mCRs 10366 while SRs 10362 includes MIS 10368 and MOS 10370.

MEMINT 20212 comprises FU 10120's interface to MEM 10112 for providing Physical Descriptors (physical addresses) to MEM 10112 to read SINs and data from and write data to MEM 10112. MEMINT 20212 includes, among other logic circuitry, MC 10226, ATC 10228, and PC 10234.

FUCTL 20214 controls fetching of SINs and data from 44 MEM 10112 and provides sequences of microinstructions for 1 10 control of FU 10120 and EU 10122 in response to SOPs.

FUCTL 20214 provides Name inputs to MC 10226 for subsequent fetching of corresponding data from MEM t J 10112. FUCTL 20214 includes, in part, MISPR 10356, RCWS 10358, Fetch Unit S-Interpreter Dispatch Table (FUSDT) Q 15 11010, and Fetch Unit S-Interpreter Table (FUSITT) 11012.

Having described the overall structure of FU 10120, in particular with regard to previous descriptions in ,Chapter 1 of this description, DESP 20210, MEMINT 20212, and FUCTL 20214 will be described in further detail below, and in that order.

H 1. Description--Ptocessof-2-2c0--fs-.

-20-2 0 1' -1 8 2 3 2 As described above, DESP 20210 comprises a 32 bit CPU for performing all usual arithmetic and logic operations on data. In addition, a primary function of DESP 20210 is generation and manipulation of entries for, for example, Name Tables (NTs) 10350, ATC 10228, and PC 10234, and generation and manipulation of logical uV -231descriptors. As previously described, with reference to CS 10110 addressing structure, logical descriptors are logical addresses, or pointers, to data stored in MEM 10112. Logical descriptors are used, for example, as architectural base pointers or microcontrol pointers in ABRs 10364 and mCRs 10366 as shown in Fig. 103, or as linkage and local pointers of Procedure Frames 10412 as shown in Fig. 104. In a further example, logical S° descriptors generated by DESP 20210 and corresponding to certain operand Names are stored in MC 10226, where they are subsequently accessed by those Names appearing in SINs fetched from MEM 10112 to provide rapid translation between operand Names and corresponding logical descriptors.

15 As has been previously discussed with reference to CS 10110 addressing structure, logical descriptors provided Sto ATU 10228, from DESP 20210 or NC 10226, are translated SR'. by ATU 10228 to physical descriptors which are actual physical addresses of corresponding data stored in MEM 10112. That data subsequently is provided to JP 10114, and in particular to FU 10120 or EU 10122, through MOD Bus 10144.

As has been previously discussed with reference to 4 MEM 10112, each data read to JP 10114 from MEM 10112 may contain up to 32 bits of information. If a particular data item referenced by a logical descriptor contains more than 32 bits of data, DESP 20210 will, as described further below, generate successive logical descriptors, .i o, -232each logical descriptor referring to 32 bits or less of information, until the entire data item has been read from MEM 10112. In this regard, it should be noted that NC 10226 may contain logical descriptors only for data items of 255 bits or less in length. All requests to MEM 10112 for data items greater than 32 bits in length are generated by DESP 20210. Most of data items operated on by CS 10110 will, however, be 32 bits or less in length so that NC 10226 is capable of handling most operand 9 e Names to logical descriptor translations.

As described above, DESP 20210 includes AONP 20216, OFFP 20218, and LENA' 20220. OFFP 20218 comprises a St f general purpose 32 bit CPU with additional logic circuitry for generating and manipulating table and cache entries, as described above, and for generating and c manipulating offset fields of AON pointers and logical St 4o descriptors. AONP 20216 and LENP 20220 comprise logic circuitry for generating and manipulating, respectively, SAON and length fields of AON pointers and logical descriptors. As indicated in Fig. 202, GRF 10354 is vertically divided in three parts. A first part resides j in ANOP 20216 and, in additon to random data, contains AON fields of logical descriptors. Second and third parts reside, respectively, in OFFP 20218 and LENP 20220 and, in addition to containing random data, respectively contain offset and length fields of logical descriptors.

AON, Offset, and length portions of GRF 10354 residing respectively in AONP 20216, OFFP 20218, and LENP 20220 *1.

L, i' i -233- .9 9.

r*

I:

CL

C C C

CCCI

I; IC p Pp..

*0 are designated, respectively, as AONGRF, OFFGRF, and LENGRF. AONGRF portion of GRF 10354 is 28 bits wide while OFFGRF and LENGRF portions of GRF 10354 are 32 bits in width. Although shown as divided vertically into three parts, GRF 10354 is addressed and operates as a unitary structure. That is, a particular address provided to GRF 10354 will address corresponding horizontal segments of each of GRF 10354's three sections residing in AONP 20216, OFFP 20218, and LENP 20220.

10 a. Offset--Processo--2-2-l-BS ture Referring first to OFFP 20218, in addition to being a 32 bit CPU and.generating and manipulating table and cache entries and offset fields of AON pointers and logical descriptors, OFFP 20218 is DESP 20210's primary path for receiving data from and transferring data to MEM 10112. OFFP 20218 includes Offset Input Select Multiplexer (OFFSEL) 20238, OFFGRF 20234, Offset Multiplexer Logic (OFFMUX) 20240, Offset ALU (OFFALU) 20242, and Offset ALU A Inputs Multiplexer (OFFALUSA) 20244.

OFFSEL 20238 has first and second 32 bit data inputs connected from, respectively, MOD Bus 10144 and JPD Bus 10142; OFFSEL 20238 has a third 32 bit data input connected from a first output of OFFALU 20242, a fourth 28 bit data input connected from a first output of AONGRF 20232, and a fifth 32 bit data input connected from OFFSET Bus 20228. OFFSEL 20238 has a first 32 bit output connected to input of OFFGRF 20234 and a second 32 bit ;4 I I Pi

V-

I n j; -234output connected to a first input of OFFMUX 20240.

OFFMUX 20240 has second and third 32 bit data inputs connected from, respectively, MOD Bus 10144 and JPD Bus 10142. OFFMUX 20240 also has a fourth 5 bit data input connected from Bias Logic (BIAS) 20246 and LENP 20220, described further below, and fifth 16 bit data input connected from NAME Bus 20224. Thirty-two bit data output of OFFGRF 20234 and first 32 bit data output of .OFFMUX 20240 are connected to, respectively, first and second data inputs of OFFALUSA 20244. A first 32 bit data output of OFFALUSA 20244 and a second 32 bit data output of OFFMUX 20240 are connected, respectively, to first and second data inputs of OFFALU 20242. A second 32 bit data output of OFFALUSA 20244 is connected to OFFSET Bus 20228. A first 32 bit data output of OFFALU v 620242 is connected to JPD Bus 10142, to a first input of AON Input Select Multiplexer (AONSEL) 20248 and AONP S20216, and, as described above, to a third input of OFFSEL 20238. A second 32 bit data output of OFFALU 20242 is connected to OFFSET Bus 20228 and third 16 bit output is connected to NAME Bus 20224.

b. AON--P~ocessor--2-2-SfteSt u-r e Referring to AONP 20216, a primary function of AONP 20216 is that of containing AON fields of AON pointers and logical descriptors. In addition, those portions of AONGRF 20232 not otherwise occupied by AON pointers and logical descriptors may be used as a 28 bit wide general register area by JP 10114. These portions of AONGRF i 235- 20232 may be so used either alone or in conjunction wih corresponding portions of OFFGRF 20234 and LENGRF 20236.

AONP 20216 includes AONSEL 20248 and AONGRF 20232. As previously described, a first 32 bit data input AONSEL 20248 is connected from a first data output of OFFALU 20242. A second 28 bit data input of AONSEL 20248 is connected from 28 bit output of AONGRF 20232 and from AON Bus 20230. A third 28 bit data input of AONSEL 20248 is 2 connected from logic zero, that is a 28 bit input wherein 10 each input bit is set to logic zero. Twenty-eight bit data output of AONSEL 20248 is connected to data input of AONGRF 20232. As just described, 28 bit data output of AONGRF 20232 is connected to second data input of AONSEL 20248, and is connected to AON Bus 20230.

c. Lenah--Ptocesso--20-2O--St-rcttu--e Referring finally to LENP 20220, a primary function of LENP 20220 is the generation manipulation of length S4 fields of AON pointers and physical descriptors. In addition, LENGRF 20236 may be used, in part, either alone or in conjunction with corresponding address spaces of AONGRF 20232 and OFFGRF 20234, as general registers for storage of data. LENP 20220 includes Length Input Select Multiplexer (LENSEL) 20250, LENGRF 20236, BIAS 20246, and Length ALU (LENALU) 20252. LENSEL 20250 has first and second data inputs connected from, respectively, LENGTH Bus 20226 and OFFSET Bus 20228. LENGTH Bus 20226 is eight data bits, zero filled while OFFSET Bus 20228 is 32 i data bits. LENSEL 20250 has a third 32 bit data input Pi L A j l -236connected from data output of LENALU 20252. Thirty-two bit data output of LENSEL 20250 is connected to data input of LENGRF 20236 and to a first data input of BIAS 20246. Second and third 32 bit data inputs of BIAS 20246 are connected from, respectively, Constant and Literal outputs of FUSITT 11012 as will be described further below. Thirty-two bits data output of LENGRF «20236 is connected to JPD Bus 10142, to ,Write Length o 1 Input (WL) input of NC 10226, and to a first input of 1 0 LENALU 20252. Five bit output of BIAS 20246 is connected S. to a second input of LENALU 20252, to LENGTH Bus 20226, and, as previously described, to a fourth input of OFFMUX 20240. Thirty-two bit output of LENALU 20252 is connected, as stated above, to third input of LENSEL 20250.

Having described the overall operation and the structure of DESP 20210, operation of DESP 20210 will be described next below in further detail.

d. Descfriptor-ProceQsr -2O l--Opeyrio-n a.a. Off set-Selector--2-8 Referring to OFFP 20218, GRF 10354 includes GR's 10360 and SR's 10362. GR's 10360 in turn contain ABR's 10364, mCR's 10366, and a set of general registers. SR's 10362 include MIS 10368 and MOS 10370. GRF 10354 is vertically divided into three parts. AONGRF 20232 is 28 bits wide and resides in AONP 20216, LENGRF 10354 is 32 bits wide and resides in LENP 20220, and OFFGRF 20234 is K 32 bits wide and resides in OFFP 20218. AONGRF 20232, l. t, flrr^ 237- OFFGRF 20234, and LENGRF 20236 may be comprised of Fairchild 93422s.

In addition to storing offset fields of AON pointers and logical descriptors, those portions of OFFGRF 20234 not reserved for ABR's 10365, mCR's 10366, and SR's 10362 may be used as general registers, alone or in conjunction with corresponding portions AONGRF 20232 and LENGRF S20236, when OFFP 20218 is being utilized as a general purpose, 32 bit CPU. OFFGRF 20234 as will be describe3 10 further below, is addressed in parallel with AONGRF 20232 and LENGRF 20236 by address inputs provided from FUCTL 20214.

S" OFFSEL 20238 is a multiplexer, comprised for example of SN74S244s'and SN74S257s, for selecting data inputs to be written into selected address locations of OFFGRF 20234. OFFSEL 20238's first data input is from MOD Bus S. t 10144 and is the primary path for data transfer between MEM 10112 and DESP 20210. As previously described, each data read from MEM 10112 to JP 10114 is a single 32 bit word where between one and 32 bits may contain actual data. If a data item to be read from MEM 10112 contains Smore than 32 bits of data, successive read operations are performed until the entire data item has been transferred.

OFFSEL 20238's second data input is from JPD Bus 10142. As will be described further below, JPD Bus 10142 is a data transfer path by which data outputs of FU 10120 and EU 10122 are written into MEM 10112. OFFSEL 20238's ki iS- -238input of JPD Bus 10142 thereby provides a wrap around path by which data present at outputs of FU 10120 or EU 10122 may be transferred back into DESP 20210 for further use. For example, as previously stated a first output of OFFALU 20242 is connected to JPD Bus 10142, thereby allowing data output of OFFP 20218 to be returned to OFFP 20218 for further processing, or to be transferred to AONP 20216 or LENP 20220 as will be described further below. In addition, output of LENGRF 20236 is also 10 connected to JPD Bus 10142 so that length fields of AON *pointers or physical descriptors, or data, may be read from LENGRF 20236 to OFFP 20218. This path may be used, j for example, when LENGRF 20236 is being used as a general purpose register for storing data or intermediate results of arithmetic or logical operations.

SvSt OFFSEL 20238's third input is provided from OFFALU 4 1,u 20242's output. This data path thereby provides a wrap around path whereby offset fields or data residing in OFFGRF 20234 may be operated on and returned to OFFGRF 20234, either in the same address location as originally read from or to a different address location. OFFP 20218 wrap around path from OFFALU 20242's output to OFFSEL S" 4 20238's third input, and thus to OFFGRF 20234, may be utilized, for example, in reading from MEM 10112 a data item containing more than 32 bits of data. As previously described, each read operation from MEM 10112 to JP 10114 A is of a 32 bit word wherein between one and 32 bits may contain actual data. Transfer of a data word containing L A r l~ i i I I L i -239more than 32 bits is accomplished by performing a succession of read operations from MEM 10112 to JP 10114. For example, if a requested data item contains bits of data, that data item will be transferred in three consecutive read operations. First and second read operations will each transfer 32 bits of data, and final read operation will transfer the remaining 6 bits of S. data. To read a data item of greater than 32 bits from MEM 10112 therefore, DESP 20210 must generate a sequence of logical descriptors, each defining a successive 32 bit S" segment of that data item. Final logical descriptor of the sequence may define a segment of less than 32 bits, for example, six bits as in the example just stated. tn each successive physical descriptor, offset field must be incremented by value of length field of the preceding ~physical descriptor to define starting addresses of successive data items segments to be transferred. Length field of succeeding physical descriptors will, in general, remain constant at 32 bits except for final transfer which may be less than 32 bits. Offset field will thereby usually be incremented by 32 bits at each Stransfer until final transfer. OFFP 20218's wrap around data path from OFFALU 20242's output to third input of OFFSEL 20238 may, as stated above, be utilized in such sequential data transfer operations to write incremented or decremented offset field of a current physical descriptor back into OFFGRF 20234 to be offset field of a next succeeding physical descriptor.

1 -240- In a further example, OFFP 20218's wrap around path from OFFALU 20242's output to third input of OFFSEL 20238 may be used in resolving Entries in Name Tables 10350, that is Name resolutions. In Name resolutions, as previously described, offset fields of AON pointers, for example Linkage Pointers 10416, are successively added and subtracted to provide a final AON pointer to a U desired data item.

OFFSEL 20238's fourth input, from AONGRF 20232's 10 output, may be used to transfer data or AON fields from S* AONGRF 20232 to OFFGRF 20234 or OFFMUX 20240. This data path may be used, for example, when OFFP 20218 is used to generate AON fields of AON pointers or physical descriptors or when performing Name evaluations.

Finally, OFFSEL 20238's fifth data input from OFFSET r r t Bus 20228 allows offset fields on OFFSET Bus 20228 to be written into OFFGRF 20234 or transferred into OFFMUX 20240. This data path may be used, for example, to copy offset fields to OFFGRF 20234 when JP 10114 is performing a Name evaluation.

SReferring now to OFFMUX 20240, OFFMUX 20240 includes logic circuitry for manipulating individual bits of 32 bit words. OFFMUX 20240 may be used, for example, to increment and decrement offset fields by length fields when performing string transfers, and to generate entries for, for example, MHT 10716 and MFT 10718. OFFMUX 20240 may also be used to aid in generating and manipulating AON, OFFSET, and LENGTH fields of physical descriptors and AON pointers.

K -241b b. Offset--Mu-tip-lexer--2-24--Detaieed Stucture--ictrj.---2 Referring to Fig. 238, a more detailed, partial block diagram of OFFMUX 20240 is shown. OFFMUX 20240 includes Offset Multiplexer Input Selector (OFFMUXIS) 23810, which for example may be comprised of SN74S373s and SN74S244s and Offset Multiplexer Register (OFFMUXR) 23812, which for example may be comprised of SN74S374s. OFFMUX 20240 also includes Field Extraction Circuit (FEXT) 23814, which may for example be comprised of SN74S257s, and SOffset Multiplexer Field Selector (OFFMUXFS) 23816, which for example may be comprised of SN74S257s and SN74S374s.

Finally, OFFMUX 20240 includes Offset Scaler (OFFSCALE) 23818, which may for example be comprised of AMD 25S10s, Offset Inter-element Spacing Encoder (OFFIESENC) 23820, which may for example be comprised of Fairchild 93427s and Offset Multiplexer Output Selector (OFFMUXOS) 23822, S. which may for example be comprised of AMD 25Ss, Fairchild 93427s, and SN74S244s.

S 20 Referring first to OFFMUX 20240's connections to other portions of OFFP 20218, OFFMUX 20240's first data input, from OFFSEL 20238, is connected to a first input of OFFMUXIS 23810. OFFMUX 20240's second input, from MOD Bus 10144, is connected to a second input of OFFMUXIS 23810. OFFMUX 20240's third input, from JPD Bus 10142, is connected to a first input of OFFMUXFS 23816 while OFFMUX 20240's fourth input, from BIAS 20246, is connected to a first input of OFFMUXOS 23822. OFFMUX 20240s fifth input, from NAME Bus 20224, is connected to !r i ,.f -242a second input of OFFMUXFS 23816. OFFMUX 20240's first output, to OFFALUSA 20244, is connected from output of OFFMUXR 23812 while OFFMUX 20240's second output, to OFFALU 20242, is connected from output of OFFMUXOS 23822.

Referring to OFFMUX 20240's internal connections, 32 bit output of OFFMUXIS 23810 is connected to input OFFMUXR 23812 and 32 bit output of OFFMUXR 23812 is connected, as described above, as first output of OFFMUX 20240, and as a third input of OFFMUXFS 23816. Thirty- 10 two bit output of OFFMUXR 23812 is also connected to "0 input of FEXT 23814. OFFMUXFS 23816's first, second and S third inputs are connected as described above. A fourth input of OFFMUXFS 23816 is a 32 bit input wherein 31 bits S.are set to logic zero and 1 bit to logic 1. A fifth input is a 32 bit input wherein 31 bits are set to logic 1 and 1 to logic 0. A sixth input of OFFMUXFS 23816 is a S32 bit literal input provided from FUSITT 11012 and is a 32 bit binary number comprising a part of a microinstruction FUCTL 20214, described below. OFFMUXFS 23816's seventh and eighth input are connected from FEXT 23814. Input 7 comprises FIU and TYPE fields of Name Table Entries which have been read into OFFMUXR 23812.

Input 8 is a general purpose input conveying bits extracted from a 32 bit word captured in OFFMUXR 23812.

Z 25 As indicated in Fig. 238, OFFMUXFS 23816's first, third, fourth, fifth, and sixth inputs are each 32 bit inputs which are divided to provide two 16 bit inputs each. I That is, each of these 32 bit inputs is divided into a L r, -243first input comprising bit 0 to 15 of that 32 bit input, and a second input comprising bits 16 to 31.

Thirty-two bit output of OFFMUXFS 23816 is connected to inputs of OFFSCALE 23818 and OFFIESENC 23820. As indicated in Fig. 238, Field Select Output (FSO) of OFFMUXFS 23816 is a 32 bit word divided into a first word including 0 to 15 and a second word including bits 16 to 31. Output FSO of OFFMUXFS 23816, as will be described further below, thereby reflects the divided structure of OFFMUXFS 23816's first, third, fourth, fifth, and sixth inputs.

Logical functions performed by OFFMUXFS 23816 in generating output FSO, and which will be described in further detail in following descriptions, include: Passing the contents of OFFMUXR 23812 directly through OFFMUXFS 23816; Passing a 32 bit word on JPD Bus 10142 directly S* through OFFMUXFS 23816; Passing a literal value comprising a part of a 20 microinstruction from FUCTL 20214 directly through OFFMUXFS 23816; S(4) Forcing FSO to be literal values 0000 0000 Forcing FSO to be literal value 0000 001 Extracting Name Table Entry fields; Accepting a 32 bit word from OFFMUXR 23812 or JPD Bus 10142, or 32 bits of a microinstruction from FUCTL 20214, and passing the lower 16 bits while forcing the upper 16 bits to logic 0; i l- t i ia..

1 j L il 1 1 1 1 1 1 1 1 -244- Accepting a 32 bit word from OFFMUXR 23812 or JPD Bus 10142, or 32 bits of miroinstruction from FUCTL 20214, and passing the higher 16 bits while forcing the lower 16 bits to logic 0; Accepting a 32 bit word from OFFMUXR 23812, or JPD Bus 10142, or Name Bus 20224, or 32 bits of a microinstruction from FUCTL 20214, and passing the lower 16 bits while sign extending bit 16 to the upper 16 bits; and, 10 (10) Accepting a 32 bit word from Name Bus 20224 and *passing the lowest 8 bits while sign extending bit 24 to Not t the highest 24 bits.

j Thirty-two bit output of OFFSCALE 23818 and 3 bit output of OFFIESENC 23820 are connected, respectively, to second and third inputs of OFFMUXOS 23822. OFFMUXOS 23822's first input is, as described above, OFFMUX t'V 20240's fourth input and is connected from output BIAS St 20246. Finally, OFFMUXOS 23822's 32 bit output, OFFMUX (0-31) is OFFMUX 20240's second output and as previously 20 described as connected to a second input of OFFALU 20242.

c. C.Offset--Mtu~ pexeT-2-2O-2-etailed Operaton t. a. aaa. Interna--Ope-a^t+on Having desctibed the structure of OFFMUX 20240 as shown in Fig. 238, operation of OFFMUX 20240 will be described below. Internal operation of OFFMUX 20240, as 4; shown in Fig. 238, will be described first, followed by I description of OFFMUX 20240's operation with regard to DESP 20210.

L

-245- Referring first to OFFMUXR 23812, OFFMUXR 23812 is a 32 bit register receiving either a 32 bit word from MOD Bus 10144, MOD or a 32 bit word received from OFFSEL 20238, OFFSEL and is selected by OFFMUXIS 23810. OFFMUXR 23812 in turn provides those selected 32 bit words from MOD Bus 10144 or OFFSEL 20238 as OFFMUX 20240's first data output to OFFALUSA 20244, as FEXT 23814's input, and as OFFMUXFS 23816's third input.

OFFMUXR 23812's 32 bit output to OFFMUXFS 23816 is 4 provided as two parallel 16 bit words designated as OFFMUXR output (OFFMUXRO) (0-15) and (16-31). As described above, OFFMUXFS 23816's output to OFFALUSA 20244 from OFFMUXR 23812 may be right shifted 16 places and the highest 16 bits zero filled.

FEXT 23814 receives OFFMUXRO (0-15) and (16-31) from OFFMUXR 23812 and extracts certain fields from those 16 bit words. In particular, FEXT 23814 extracts FIU and TYPE fields from NT 10350 Entries which have been transferred into OFFMUXR 23812. FEXT 23814 may then 20 provide those FIU and TYPE fields as OFFMUXFS 23816's seventh input. FEXT 23814 may, selectively, extract certain other fields from 32 bit words residing in OFFMUXR 23812 and provide those fields as OFFMUXFS :'"23816's eighth input.

OFFMUXFS 23816 operates as a multiplexer to select certain fields from OFFMUXFS 23816's eight inputs and provide corresponding 32 bit output words, Field Select Output (FSO), comprised of those selected fields from i '1 -246- OFFMUXFS 23816's inputs. As previously described, FSO is comprised of 2, parallel 16 bit words, FSO (0-15) and FSO (16-31). Correspondingly, OFFMUX 20240's third input, from JPD Bus 10142, is a 32 bit input presented as two 16 bit words, JPD (0-15) and JPD (16-31). Similarly, OFFMUXFS 23816's fourth, fifth, and sixth inputs are each presented as 32 bit words comprised of 2, parallel 16 bit words, respectively, (0-15) and (16-31), (0-15) and (16-31), and L (0-15) and (16-31). OFFMUXFS 23816's 10 second input, from NAME Bus 20224, is presented as a single 16 bit word, NAME (16-31), while OFFMUXFS 23816's t t inputs from FEXT 23814 are each less than 16 bits in St width. OFFMUXFS 23816 may, for a single 32 bit output SJ word, select FSO (0-15) to contain one of corresponding 16 bit inputs JPD or L Similarly, FSO (16-31) of that 32 bit output t word may be selected to contain one of NAME (16-31), JPD (16-31), 0 (16-31), 1 (16-31), L (16-31), or inputs 7 and 4 £4 8 from FEXT 23814. OFFMUXFS 23816 therefore allows 32 V 20 bit words, comprised of two 16 bit fields, to be generated from selected portions of OFFMUXFS 23816's inputs.

C OFFMUXFS 23816 32 bit output is provided as inputs to OFFSCALE 23818 and OFFIESENC 23820. Referring first to OFFIESENC 23820, OFFIESENC 23820 is used, in particular, 4 in resolving, or evaluating, NT 10350 Entries (NTEs) referring to arrays of data words. As indicted in Fig.

108, word D of an NTE contains certain information relating to inter-element spacing (IES) of data words of 1" -247an array. Word D of an NTE may be read from MEM 10112 to MOD Bus 10144 and through OFFMUX 20240 to input of OFFIESENC 23820. OFFIESENC 23820 then examines word D's IES field to determine whether inter-element spacing of that array is a binary multiple, that is 1, 2, 4, 8, 16, 32, or 64 bits. In particular, OFFIESENC 23820 determines whether 32 bit word D contains logic zeros in the most significant 25 bits and a single logic one in the least significant 7 bits. If inter-element spacing I 10 is such a binary multiple, starting addresses of data words of that array may be determined by left shifting of index (IES) to obtain offset fields of physical addresses I of words in the array and a slower and more complex multiplication operation is not required. In such cases, OFFIESENC generates a first output, IES Encodeable (IESENC) to FUCTL 20214 to indicate that inter-element spacing may be determined by simple left shifting.

.OFFIESENC 23820 then generates encoded output, Encoded IES (ENCIES), to OFFMUXOS 23822. ENCIES is then a coded S" 20 value specifying the amount of left shift necessary to translate index (IES) value into offsets of words in that array. As indicated in Fig. 238, ENCIES is OFFMUXOS '23822's third input.

t OFFSCALE 23818 is a left shift shift network with zero fill of least significant bits, as bits are left shifted. Amount of shift by OFFSCALE 23818 is selectable between zero and 7 bits. Thirty-two bit words transferred into OFFSCALE 23818 from OFFSCALE 23818 from -248- OFFMUXFS 23816 may therefore be left shifted, bit by bit, to selectively reposition bits within that 32 bit input word. In conjunction with OFFMUXFS 23816, and a wrap around connection provided by OFFALU 20242's output to OFFSEL 20238, OFFSCALE 23818 may be used to generate and manipulate, for example, entries for MHT 10716, MFT 10718, AOT 10712, and AST 10914, and other CS 10110 data structures.

OFFMUXOS 23822 is a multiplexer having first, second, 0 and third inputs from, respectively, BIAS 20246, OFFSCALE 23818, OFFIESENC 23820. OFFMUXOS 23822 may select any Sone of these inputs as OFFMUX 20240's second output, OFFMUX As previously described, OFFMUX 20240's i' second output is connected to a second input of OFFALU 20242.

Having described internal of OFFMUX 20240, operation

S

i of OFFMUX 20240 with regard to overall operation of DESP 20210 will be described next below.

b.b.b. Qperation-Re-la-ive-to Descr iptor-e-Pocessor -20-2-10 j OFFMUX 20240's first input, from OFFSEL 20238, allows inputs to OFFSEL to be transferred through OFFMUXIS 23810 and into OFFMUXR 23812. This input allows OFFMUXR 23812 to be loaded, under control of FUCTL 20214 microinstructions, with any input of OFFSEL 20238. In a particular example, OFFALU 20242's output may be fed back through OFFSEL 20238's third input and OFFMUX 20240's i 1 1 k( j -249first input to allow OFFMUX 20240 and OFFALU 20242 to perform reiterative operations on a single 32 bit word.

OFFMUX 20240's second input, from MOD Bus 10144, allows OFFMUXR 23812 to be loaded directly from MOD Bus 10144. For example, NTEs from a currently active procedure may be loaded into OFFMUXR 23812 to be operated upon as described above. In addition, OFFMUX 20240's second input may be used in conjunctionwith OFFSEL 20238's first input, from MOD Bus 10144, as parallel input paths to OFFP 20218. These parallel input paths allow pipelining of OFFP 20218 operations by allowing \OFFSEL 20238 and OFFGRF 20234 to operate independently from OFFMUX 20240. For example, FU 10120 may initiate a read operation from MEM 10112 to OFFMUXR 23812 during a first microinstruction. The data so requested will appear on MOD Bus 10144 during a second microinstruction and may be loaded into OFFMUXR 23812 through OFFMUX 20240's second input from MOD Bus 10144. Concurrently, FU 10120 may initiate, at start of second microinstruction, an independent operation to be performed by OFFSEL 20238 and OFFGRF 20234, for example loading output of OFFALU 20242 into OFFGRF 20234.

r Therefore, by providing an independent path into OFFMUX 20240 from MOD Bus 10144, OFFSEL 20238 is free to perform other, concurrent data transfer operations while a data transfer from MOD Bus 10144 to OFFMUX 20240 is being performed.

I' 1 ~i a 1 i -250- OFFMUX 20240's third input, from JPD Bus 10142, is a general purpose data transfer path. For example, data from LENGRF 20236 or OFFALU 20242 may be transferred into OFFMUX 20240 through JPD Bus 10142 and OFFMUX 20240's third input.

OFFMUX 20240's fourth input is connected from BIAS 20246 and primarily used during string transfers as described above. That is, length fields of physical descriptors generated for a string transfer may be 10 transferred into OFFMUX 20240 through OFFMUX 20240's fourth input to increment or decrement, offset fields of those physical descriptors in OFFALU 20242.

OFFMUX 20240's fifth input is connected from NAME Bus 20224. As will be described further below, Names are provided to NC 10226 by FUCTL 20214 to call, from MC 10226, logical descriptors corresponding to Names appearing on MOD Bus 10144 as part of sequences of SINs.

As each Name is presented to NC 10226, that Name is transferred into and captured in Name Trap (NT) 20254.

Upon occurrence of an NC 10226 miss, that is NC 10226 does not contain an entry corresponding to a particular Name, that Name is subsequently transferred from NT 20254 to OFFMUX 20240 through NAME Bus 20224 and OFFMUX 20240's fifth input. That Name, which is previously described as an 8, 12, or 16 bit binary number, may then be scaled, that is multiplied by a NTE size. That scaled Name may then be added to Name Table Pointer (NTP) from mCRs 10366 to obtain the address of a corresponding NTE in an NT I, Ax* -L -251a C4 #4 t r *t t

I:

r: 4r 44 4 44 9* 4 4 4*4e 4*i 4. 9 4*4944 9 4 10350. In addition, a Name resulting in a NC 10226 miss, or a page fault in the corresponding NT 10350, or requiring a sequence of Name resolves, may be transferred into OFFGRF 20234 from OFFMUX 20240, through OFFALU 20242 and OFFSEL 20238 third input. That Name may subsequently be read, or restored, from OFFGRF 20234 as required.

Referring now to outputs of OFFMUX 20240, OFFMUX 20240's first output, from OFFMUXR 23812, allows contents of OFFMUXR 23812 to be transferred to first input of OFFALU 20242 through OFFALUSA 20244. OFFMUX 20240's second output, from OFFMUXOS 23822, is provided directly to second input of OFFALU 20242. OFFALU 20242 may be concurrently provided with a first input from OFFMUXR 23812 and a second input, for example a manipulated 15 offset field, from OFFMUXOS 23822.

Referring to OFFALUSA 20244, OFFALUSA 20244 is a multiplexer. OFFALUSA 20244 may select either output of OFFGRF 20234 or first output of OFFMUX 20240 to be either first input of OFFALU 20242 or to be OFFP 20218's output to OFFSET Bus 20228. For example, an offset field from OFFGRF 20234 may be read to OFFSET Bus 20228 to comprise offset field of a current logical descriptor, and concurrently read into OFFALU 20242 to be incremented or decremented to generate offset field of a subsequent logical descriptor in a string transfer.

're tl r i

~I

Bi 1 -252- OFFALU 20242 is a general purpose, 32 bit arithmetic and logic unit capable of performing all usual ALU operations. For example, OFFALU 20242 may add, subtract, increment, or decrement offset fields of logical descriptors. In addition, OFFALU 20242 may serve as a transfer path for data, that is OFFALU 20242 may transfer input data to OFFALU 20242's outputs without operating upon that data. OFFALU 20242's first output, as 9 described above, is connected to JPD Bus 10142, to third 10 input of OFFSEL 20238, and to first input of AONSEL 20248. Data transferred or manipulated by OFFALU 20242 may therefore be transferred on to JPD Bus 10142, or wrapped around into OFFP 20218 through OFFSEL 20238 for subsequent or reiterative operations. OFFALU 20242's 15 output to AONSEL 20248 may be used, for example, to load SAON fields of AON pointers or physical descriptors generated by OFFP 20218 into AONGRF 20232. In addition, this data path allows FU 10120 to utilize AONGRF 20232 as, for example, a buffer or temporary memory space for intermediate or final results of FU 10120 operations.

OFFALU 20242's output to OFFSET Bus 20228 allows logical descriptor offset fields to be transferred onto OFFSET Bus 20228 directly from OFFALU 20242. For example, a logical descriptor offset field may be generated by OFFALU 20242 during a first clock cycle, and 9 transferred immediately onto OFFSET Bus 20228 during a second clock cycle.

*I

"U -253- OFFALU 20242's third output is to NAME Bus 20224. As will be described further below, NAME Bus 20224 is address input (ADR) to NC 10226. OFFALU 20242's output to NAME Bus 20224 thereby allows OFFP 20218 to generate or provide addresses, that is Names, to NC 10226.

Having described operation of OFFP 20218, operation of LENP 20220 will be described next below.

e. Length--Processor-20-2-2-f i--2-399 Referring to Fig. 202, a primary function of LENP 10 20220 is generation and manipulation of logical descriptor length fields, including length fields of logical descriptors generated in string transfers. LENP 20220 includes LENGRF 20236, LENSEL 20250, BIAS 20246, and LENALU 20252. LENGRF 20236 may be comprised, for example, of Fairchild 93422s. LENSEL 20250 may be comprised of, for example, SN74S257s, SN74S157s, and SN74S244s, and LENALU 20252 may be comprised of, for example, SN74S381s.

As previously described, LENGRF 20236 is a 32 bit wide vertical section of GRF 10354. LENGRF 20236 operates in parallel with OFFGRF 20234 and AONGRF 20232 and contains, in part, length fields of logical descriptors. In addition, also as previously described, LENGRF 20236 may contain data.

LENSEL 20250 is a multiplexer having three inputs and providing outputs to LENGRF 20236 and first input of BIAS 20246. LENSEL 20250's first input is from Length Bus 20226 and may be used to write physical descriptor or

II

-254length fields from LENGTH Bus 20226 into LENGRF 20236 or into BIAS 20246. Such length fields may be written from LENGTH Bus 20226 to LENGRF 20236, for example, during Name evaluation or resolve operations. LENSEL 20250's second input is from OFFSET Bus 20228. LENSEL 20250's second input may be used, for example, to load length fields generated by OFFP 20218 into LENGRF 20236. In addition, data operated upon by OFFP 20218 may be read •r into LENGRF 20236 for storage through LENSEL 20250's 10 second input.

LENSEL 202501s third input is from output of LENALU t 20252 and is a wrap around path to return output of S< LENALU 20252 to LENGRF 20256. LENSEL 20250's third input may, for example, be used during string transfers when length fields of a particular logical descriptor is f incremented or decremented by LENALU 20252 and returned 9 to LENGRF 20236. This data path may also, for example, L OA be used in moving a 32 bit word from one location in l s LENGRF 20236 to another location in LENGRF 20236. As stated above, LENSEL 20250's output is also provided to first input BIAS and allows data appearing at first, S' second, or .third inputs of LENSEL 20250 to be provided to first input of BIAS 20246.

BIAS 20246, as will'be described in further detail below, generates logical descriptor length fields during string transfers. As described above, no more than 32 bits of data may be read from MEM 10112 during a single read operation. A data item of greater than 32 bits in i -255length must therefore be transferred in a series, or string, of read operations, each read operation transferring 32 bits or less of data. String transfer logical descriptor length fields generated BIAS 20246 are provided to LENGTH Bus 20226, to LENALU 20252 second input, and to OFFMUX 20240's fourth input, as previously described. These string transfer logical descriptor length fields, referred to as bias fields are provided to LENGTH Bus 20226 by BIAS 20246 to be length fields of the 10 series of logical descriptors generated by DESP 20210 to 4 execute a string transfer. These bias fields are provided to fourth input OFFMUX 20240 to increment or decrement offset fields of those logical descriptors, as.

previously described. These bias fields are provided to second input of LENALU 20252, during string transfers, to correspondingly decrement the length field of a data item Sbeing read to MEM 10112 in a string transfer. BIAS 20246 will be described in greater detail below, after LENALU 20252 is first briefly described.

a.a. nenath-Ahf 20-252 LENALU 20252 is a general purpose, 32 bit arithmetic and logic unit capable of executing all customary arithmetic and logic operations. In particular, during a string transfer of a particular data item LENALU 20252 receives that data items length field from LENGRF 20236 and successive bias fields from BIAS 20246. LENALU 20252 then decrements that logical descriptor's current length field to generate length field to be used during next read operation of the string transfer, and transfers new length field back into LENGRF 20236 through LENSEL 20250's third input.

|T I -256b.b. BI-AS-2 Referring to Fig, 239, a partial block diagram of BIAS 20246 is shown. BIAS 20246 includes Bias Memory (BIASM) 23910, Length Detector (LDET) 23912, Next Zero Detector (NXTZRO) 23914, and Select Bias (SBIAS) 23916.

Input of LDET 23912 is first input of BIAS 20246 and connected from output of LENSEL 20250. Output of LDET 23912 is connected to data input of BIASM 23910, and data Soutput of BIASM 23910 is connected to input of NXTZRO 23914. Output of NXTZRO 23914 is connected to a first input of SBIAS 23916. A second input of SBIAS 23916 is BIAS 20246's second input, L8, and is connected from an output of FUCTL 20214. A third input of SBIAS 23916 is BIAS 20246's third input, L, and is connected from yet another output of FUCTL 20214. Output of SBIAS 23916 is output of BIAS 20246 and, as described above, is connected to LENGTH Bus 20226, to a second input of LENALU 20252, and to fourth input of OFFMUX 20240.

BIASM 23910 is a 7 bit wide random access memory having a length equal to, and operating and addressed in parallel with, SR's 10362 of GRF 10354. BIASM 23910 has an address location corresponding to each address location of SR's 10362 and is addrestsed concurrently with those address locations in SR's 10362. BIASM 23910 may be comprised, for example, of AMD 27S03As.

BIASM 23910 contains a bias value of each logical descriptor residing in SR's 10362. As described above, a bias value is a number representing number of bi'% to be -257read from MEM 10112 in a particular read operation when a data item having a corresponding logical descriptor, with a length field stored LENGRF 20236, is to be read from MEM 10112. Initially, bias values are written into BIASM 23910, in a manner described below, when their corresponding length fields are written into LENGRF 20236. If a particular data item has a length of less S. than 32 bits, that data item's initial bais value will represent that data items actual length. For example, if a data item has a length of 24 bits the associated bias value will be a 6 bit binary number representing 24.

That data item's length field in LENGRF 20236 will similarly contain a length value of 24. If a particular item has a length of greater than 32 bits for example, S 15 bits as described in a previous example, that data item a must be read from MEM 10112 in a string transfer operation. As previously described, a string transfer is a scries of read operations transferring 32 bits at a time from MEM 10112, with a final transfer of 32 bits or less completing transfer of that data item. Such a data item's initial length field entry in LENGRF 20236 will too*# contain, using the same example as previously described, a value of 70. That data item's initial bais entry written into a corresponding address space of BIASM 23910 will contain a bias value of 32. That initial bias value Sof 32 indicates that at least the first read operation required to transfer that data item from MEM 10112 will transfer 32 bits of data.

-258- When a data item having a length of less than 32 bits, for example 24 bits, is to be read from MEM 10112, that data item's bias value of 24 is read from BIASM 23910 and provided to LENGTH Bus 20226 as length field of logical descriptor for that read operation.

Concurrently, that bias value of 24 is subtracted from that data items length field read from LENGRF 20236.

Subtracting that bias value from that length value will yield a result of zero, indicating that no further read operations are required to complete transfer of that data item.

If a data item having, for example, a length of bits is to be read from MEM 10112, that data item's initial bias value of 32 is read from BIASM 23910 to LENGTH Bus 20226 as length field of first logical descriptor of a string transfer. Concurrently, that data item's initial length field is read from LENGRF 20236.

That data item's initial bias value, 32, is subtracted from that data item's initial length value, 70, and LENALU 20252. The result of that subtraction operation 4 is the remaining length of data item to be transferred in one or more subsequent read operations. In this example, subtracting initial bias value from initial length value indicates that 38 bits of that data item remain to be transferred. LENALU 20252's output representing results of this subtraction, for example 38, are transferred to LENSEL 20250's third input to LENGRF 20236 and written into address location from which that data item's initial L ti -259length value was read. This new length field entry then represents remaining length of that data item.

Concurrently, LDET 23912 examines that residual length value being written into LENGRF 20236 to determine whether remaining length of that data item is greater than 32 bits or is equal to or less than 32 bits. If remaining length is greater than 32 bits, LDET 23912 generates a next bias value of 32, which is written into BIASM 23910 and same address location that held initial bias value. If remaining data item length is less than 32 bits, LDET 23912 generates a 6 bit binary number representing actual remaining length of data item to be transferred. Actual remaining length would then, again, be written into BIASM 23910 address location originally containing initial bias value. These operations are also performed by LDET 23912 in examining initial length field 4 *f and generating a corresponding initial bias value. These read operations are continued as described above until LDET 23912 detects that remaining length field is 32 bits or less, and thus that transfer of that data item will be OW:N. completed upon next read operation. When this event is detected, LDET 23912 generates a seventh bit input into S BIASM 23910, which is written into BIASM 23910 together with last bias value of that string transfer, indicating that remaining length will be zero after next read operation. When a final bias value is read from BIASM 23910 at start of next read operation of that string transfer, that seventh bit is examined by NXTZRO 23914 which subsequently generates a test condition output, Last Read (LSTRD) to FUCTL 20214, FUCTL 20214 may then -260terminate execution of that string transfer after that last read operation, if the transfer has been successfully completed.

As previously described, the basic unit of length of a data item in CS 10110 is 32 bits. Accordingly, data items of 32 or fewer bits may be transferred directly while data items of more than 32 bits require a string transfer. In addition, transfer of a data item through a 0 o, string transfer requires tracking of the transferred 10 length, and remaining length to be transferred, of both w the data item itself and the data storage space of the 4 A location the data item is being transferred to. As such, BIAS 20246 will store, and operate with, in the manner described above, length and bias fields of the logical descriptors of both the data item and the location the data item is being transferred to. FUCTL 20214 will receive an LSTRD test condition if bias field of source 4descriptor becomes zero before or concurrently with that 'of the destination, that is a completed transfer, or if bias field of destination becomes zero before that of the source, and may provide an appropriate microcode control 4 response. It should be noted that if source bias field 4 •.•becomes zero before that of the destination, the remainder of the location that this data item is being transferred to will be filled and padded with zeros. If the data item is larger than the destination storage capacity, the destination location will be filled to capacity and FUCTL 20214 notified to initiate appropriate action.

4 -261- In addition to allowing data item transfers which are insensitive to data item length, BIAS 20246 allows string transfers to be accomplished by short, tight microcode loops which are insensitive to data item length. A string transfer, for example, from location A to location B is encoded as: Fetch from A, subtract length from bias A, and update offset and length of a; and, Store to B, subtract length from bias B, and branch to if length of B does not go to zero or fall Sthrough (end transfer) if length of B goes to zero.

Source length need not be texted as the microcode •loop continues until length of B goes to zero; as described above, B will be filled and padded with zeros if length of A is less than length of B, or B will be filled and the string transfer ended if length of A is greater than or equal to length of B.

i. LDET 23912 and NXTZRO 23914 thereby allow FUCTL 20214 to automatically initiate a string transfer upon occurrence of a single microinstruction from FUCTL 20214 initiating a read operation by DESP 20210. That 1 microinstruction initiating a read operation will then be automatically repeated until LSTRD to FUCTL 20214 from NXTZRO 23914 indicates that the string transfer is completed. LDET 23912 and NXTZRO 23914 may, respectively, be comprised for example of S74S260s, SN74S133s, SN74S51s, SN74S00s, SN74S00s, SN74S04s, SN74S02s, and SN74S32s.

-262- Referring finally to SBIAS 23916, SBIAS 23916 is a multiplexer comprised, for example, of SN74S288s, SN74S374s, and SN74S244s. SBIAS 23916, under microinstruction control from FUCTL 20214, selects BIAS 20246's output to be one of a bias value from BIASM 23910, L8, or L. SBIAS 23916's first input, from BIASM 23910, has been described above. SBIAS 23916's second input, L8, is provided from FUCTL 20214 ,and is 8 bits of S, a microinstruction provided from FUSITT 11012. SBIAS 23916's second input allows microcode selection of bias S. values to be used in manipulation of length and offset fields of logical descriptors by LENALU 20252 and OFFALU 20242, and for generating entries to MC 10226. SBIAS 23916's third input, L, is similarly provided from FUCTL 20214 and is a decoded length value derived from portions of microinstructions in FUSITT 11012. These microcode length values represent certain commonly occurring data item lengths, for example length of 1, 2, 4, 8, 16, 32, I and 64 bits. An L input representing a length of 8 bits, may be used for example in reading data from MEM 10112 on a byte by byte basis.

Having described operation of LENP 20220, operation of AONP 20216 will be described next below.

f. AON-Processor--20-2-16 a.a. AONGRP-20-232 As described above, AONP 20216 includes AONSEL 20248 and AONGRF 20232, AONGRF 20232 is a 28 bit wide vertical Ssection of GRF 10354 and stores AON fields of AON pointers and logical descriptors. AONSEL 20248 is a multiplexer for selecting inputs to be written into "i 14 -263- AONGRF 20232. AONSEL 20248 may be comprised, for example of SN74S257s. AONGRF 20232 may be comprised of, for example, Fairchild 93422s.

As previously described, AONGRF 20232's output is connected onto AON Bus 20230 to allow AON fields of AON pointers and logical descriptors to'be transferred onto AON Bus 20230 from AONGRF 20232. AONGRF 20232's output, together with a bi-directional input from AON Bus 20230, is connected to a second input of AONSEL 20248 and to a fourth input of AONSEL 20238. This data path allows AON fields, either from AONGRF 20232 or from AON Bus 20230, to be written into AONGRF 20232 or AONGRF 20234, or provided as an input to OFFMUX.20240.

b. b AON-Slecto--G}-248 AONSEL 20248's first input is, as previously described, connected from output of OFFALU 20242 and is used, for example, to allow AON fields generated or manipulated by OFFP 20218 to be written into AONGRF 20232. AONSEL 20248's third input is a 28 bit word wherein each bit is a logical zero. AONSEL 20248's third input allows AON fields of all zeros to be written into t AONGRF 20232. An AON field of all zeros is reserved to indicate that corresponding entries in OFFGRF 20234 and LENGRF 20236 are neither AON pointers nor logical descriptors. AON fields of all zeros are thereby reserved to indicate that corresponding entries in OFFGRF 20234 and LENGRF 20236 cont.ain data.

In summary, as described above, DESP 20210 includes AONP 20216, 0OFP 20218, and LENP 20220. OFFP 20218 contains a vertical section of GRF 10354, OFFGRF 20234, 2 -264- 4* I IO SI Ir for storing offset fields of AON pointers and logical descriptors, and for containing data to be operated upon by DESP 20210. OFFP 20218 is principal path for transfer of data from MEM 10112 to JP 10114 and is a general purpose 32 bit arithmetic and logic unit for performing all usual arithmetic and logic operations. In addition, OFFP 20218 includes circuitry, for example OFFMUX 20240, for generation and manipulation of AON, OFFSET, and LENGTH fields of logical descriptors and AON pointers.

10 OFFP'20218 may also generate and manipulate entries for, for example, NC 10226, ATU 10228, PC 10234, AOT 10712, MHT 10716, MFT 10718, and other data and address structures residing in MEM 10112. LENP 20220 includes a vertical section of GRF 10354, LENGRF 20236, for storing length fields of logical descriptors, and for storing data. LENP 20220 further includes BIAS 20246, used in conjunction with LENGRF 20236 and LENALU 20252, for providing length fields of logical descriptors for MEM 10112 read operations and in particular automatically performing string transfers. AONP 20216 similarly includes a vertical section of GRF 10354, AONGRF 20232.

A primary function AONGRF 20232 is storing and providing AON fields of AON pointers and logical descriptors.

Having described structure and operation of DESP 20210, structure and operation of Memory Interface (MEMINT) 20212 will be described next below.

2. Memory-- ntfe MEMINT 20212 comrises FU 10120's interface to MEM 10112. As described above, MEMINT 20212 includes Name il 4j -265- Cache (NC) 10226, Address Translation Unit (ATU) 10228, and Protection Cache (PC) 10234, all of which have been previously briefly described. MEMINT 20212 further includes Descriptor Trap (DEST) 20256 and Data Trap (DAT) 20258. Functions performed by MEMINT 20212 includes (1) resolution of Names to logical descriptors, by NC 10226; translation of logical descriptors to physical descriptors, by ATU 10228; and confirmation of access writes to objects, by PC 10234.

S 10 As shown in Fig. 202, NC 10226 adress input (ADR) is r connected from NAME Bus 20224. NC 10226 Write Length I Field Input (WL) is connected from LENGRF 20236's output. NC 10226's Write Offset Field Input (WO) and Write AON Field Input (WA) are connected, respectively, from OFFSET Bus 20228 and AON Bus 20230. NC 10226 Read AON Field Iead Offset Field and Read Length Field (RL) outputs are connected, respectively, to AON Bus 20230, OFFSET Bus 20228, and LENGTH Bus 20226.

DEST 20256's bi-directional AON (AON), Offset (OFF), and Length (LEN) ports are connected by bi-directional buses to and from, respectively, AON Bus 20230, OFFSET Bus 20228, and LENGTH Bus 20226.

PC 10234 has AON (AON) and Offset (OFF) inputs connected from, respectively, AON Bus 20230 and OFFSET Bus 20228. PC 10234 has a Write Entry (WEN) input connected from JPD Bus 10142. ATU 10228 has AON (AON) Offset (OFF) and Length (LEN) inputs connected from, respectively, AON Bus 20230, OFFSET Bus 20228, and LENGTH *1 7-K~ 7 -266ti SIt I I I t If P I rS Bus 20226. ATU 10228's output is connected to Physical Descriptor (PD) Bus 10146.

Finally, DAT 20258 has a bi-directional port connected to and from JPD Bus 10142.

a.a. Descriptor-Trap-2--2 256--and-Data Trap-20258 Referring first to DST 20256 and DAT 20258, DST 20256 is a register for receiving and capturing logical descriptors appearing on AON Bus 20230, OFFSET Bus 20228, 10 and Length Bus 20226. Similarly, DAT 20258 is a register for receiving and capturing data words appearing on JPD Bus 10142. DST 20256 and DAT 20258 may subsequently return captured logical descriptors or data words to, respectively, AON Bus 20230, OFFSET Bus 20228, and LENGTH Bus 20226, and to JPD Bus 10142.

As previously described, many CS 10110 operations, in particular MEM 10112 and JP 10114 operations, are pipelined. That is, operations are overlapped with certain sets within two or more operations being executed 20 concurrently. For example, FU 10120 may submit read request to MEM 10112 and, while MEM 10112 is accepting and servicing that request, submit a second read request. DEST 20256 and DAT 20258 assist in execution of overplapping operations by providing a temporary record of these operations. For example, a part of a read or write request to MEM 10112 by FU 10120 is a logical descriptor provided to ATU 10228. If, for example the first red request just referred to results in a ATU 10228 cache miss or a protection violation, the logical descriptor of that first request must be recovered for 1141 4 4* -267subsequent action by CS 10110 as previously described.

That logical descriptor will have been captured and stored in DEST 20256 and thus is immediately available, so that DESP 20210 is not required to regenerate that descriptor. DAT 20258 serves a similar purpose with regard to data being written into MEM 10112 from JP 10114. That is, DAT 20258 receives and captures a copy of each 32 bit word transferred onto JPD Bus 10142 by JP 10114. In event of MEM 10112 being unable to accept a 10 write request, that data may be subsequently reprovided from DAT 20258.

b.b. Name-'CacheQ-226--Addes-s Tans-ation-Utnit-+&t-2-2,-esnd Protection--eache-29-4-4Pe--K--6 Referring to NC 10226, ATU 10228, and PC 10234, these elements of MEMINT 20212 are primarily cache mechanisms c to enhance the speed of FU 10120's interface to MEM 10112, and consequently of CS 10110's operation. As described previously, NC 10226 contains a set of logical descriptors corresponding to certain operand names currently appearing in a process being executed by CS 10110. NC 10226 thus effectively provides high speed resolution of certain operand names to corresponding logical descriptors. As described above with reference to string transfers, NC 10226 will generally contain logical descriptors only for data items of less than 256 bits length. NC 10226 read and write addresses are names provided on NAME Bus 20224. Name read and write -268addresses may be provided from DESP 20210, and in particular from OFFP 20218 as previously described, or from FUCTL 20214 as will be described in a following description of FUCTL 20214. Logical descriptors comprising NC 10226 entries, each entry comprising an AON field, an Offset field, a Length field, are written into NC 10226 through NC 10226 inputs WA, WO, and WL from, respectively, AON Bus 20230, OFFSET Bus 20228, and LENGRF 20236's output. Logical descriptors read from NC 10226 in response to names provided to NC 10226 ADR input are provided to AON Bus 20230, OFFSET Bus 20228, and LENGTH Bus 20226 from, respectively, NC 10226 outputs RA, RO, and RL.

ATU 10228 is similarly a cache mechanism for providing high speed translation of logical to physical descriptors. In general, ATU 10228 will contain, at any given time, a set of logical to physical page number mappings for MEM 10112 read and write requests which are currently being made, or anticipated to be made, to MEM 20 10112 by JP 10114. As previously described, each physical descriptor is comprised of a Frame Number (FN) field, and Offset iLhin Frame (0 fields, and a Length field. As discussed with reference to string transfers, Sa physical descriptor length field, as in a logical descriptor length field, specify a data item of less than or equal to 32 bits length. Referring to ?ig. 106C, as previously discussed a logical descriptor comprised of a 14 bit AON field, a 32 bit Offset field, and Length field, wherein 32 bit logical descriptor Offset field is divided into a 18 bit Page Number field and a 14 bit I:-L i -269- Offset within Page (0 field. In translating a logicl into a physical descriptor, logical descriptor Length and 0 fields are used directly, as respectively, physical descriptor length and 0 fields. Logical descriptor AON and P fields are translated into physical descriptor FN field. Because no actual translation is required, ATU 10228 may provide logical descriptor L field and corresponding 0 field directly, that is without delay, to MEM 10112 as corresponding physical descriptor 0 and 10 Length fields. ATU 10228 cache entries are thereby comprised of physical descriptor FN fields corresponding to AON and P fields of those logical descriptors for which ATU 10228 has corresponding entries. Because physical descriptor FN fields are provided from ATU S. 15 10228's cache, rather than directly as in physical descriptor O and Length fields, a physical descriptor's I FN field will be provided to MEM 10112, for example, one clock cycle later than that physical descriptors 0 and Length fields, as has been previously discussed.

S' 20 Referring to Fig. 202, physical descriptor FN fields to be written into ATU 10228 are, in general, generated by DESP 20210. FN fields to be written into ATU 10228 are provided to ATU 10228 Data Input (DI) through JPD Bus 10142. ATU 10228 read and write addresses are comprised of AON and P fields of logical descriptors and are provided to ATU 10228's AON and OFF inputs from, respectively, AON Bus 20230 and OFFSET Bus 20228. ATU 10228 read and write addresses may be provided from DESP 20210 or, as described further below, from FUCTL 20214.

ATU 10228 FN outputs, together with 0 and Length fields -270comprising a physical descriptor, are provided to PD Bus 10146.

PC 10234 is a cache mechanism for confirming active procedure's access rights to objects identified by logical descriptors generated as a part of JP 10114 read or write requests to MEM 10112. As previously described access rights to objects are arbitrated on the basis of Ssubjects. A subject has been defined as a particular Scombination of a principal, process, and domain. A r10 principal, process, and domain are each identified by p corresponding UIDs. Each subject having access rights to an object is assigned an Active Subject Number (ASN) as described in a previous description of CS 10110's S. Protection Mechanism. The ASN of a subject currently S active in CS 10110 is stored in ASN Register 10916 in FU 10120. Access rights of a currently active subject to currently active objects are read from those objects Access Control Lists (ACL) 10918 and stored in PC 10234.

If the current ASN changes, PC 10234 is flushed of 20 corresponding access right entries and new entries, corresponding to the new ASN, are written into PC 10234.

The access rights of a particular current ASN to a particular object may be determined by indexing, or addressing, PC 10234 with the AON identifying that object. Addresses to write entries into or read entries from PC 10234 are provided to PC 10234 AON input from AON Bus 20230. Entries to be written into PC 10234 are provided to PC 10234's WEN input from JPD Bus 10142. PC 10234 is also provided with inputs, not shown in Fig. 202 for purposes of clarity, from FUCTL 20214 indicating the i1^ 2 5 ojc.Adesst rt etisit rra nre -271current operation to be perfomed by JP 10114 with respect to an object being presently addressed by FU 10120.

Whenever FU 10120 submits a read or write request concening a particular object to MEM 10112, AON field of that request is provided as an addess to PC 10234.

Access rights of the current active subject to that object are read from corresponding PC 10234 entry and compared to FUCTL 20214 inputs indicating the particular operation to be performed by JP 10114 with respect to 10 that object. The operation to be performed by JP 10114 is then compared to that active subject's access rights to that object and PC 10234 provides an output indicating whether that active subject possesses the rights required *to perform the intended operation. Indexing of PC 10234 and comparison of access rights to intended operation is performed concurrently with translation of the memory request logical descriptor to a corresponding physical descriptor by ATU 10228. If PC 10234 indicates that that active subject has the required access rights, the intended operation is executed by JP 10114. If PC 10234 :t rindicates that that active subject does not have the required access rights, PC 10234 indicates that a protection mechanism violation has occurred and interrupts execution of the intended operation.

25 c.c. Sructure-and--pepiona eneraJ-Tzed -ahe-a nd-e,- Having described overall structure and operation of NC 10226, ATU 10228, and PC 10234, structure and operation of these caches will be described in further -272detail below. Structure and operation of NC 10226, ATU 10228, and PC 10234 are similar, except that NC 10226 is a four-way set associative cache, ATU 10228 is a threeway set associative cache and PC 10234 is a two-way set associative cache.

As such, the structure and operation of NC 10226, ATU 10228, and PC 10234 will be described by reference to and description of a generalized cache similar but not necessarily identical to each of NC 10226, ATU 10228, and PC 10234. Reference will be made to NC 10226 in the description of a generalized cache next below, both to further illustrate structure and operation of the generalized cache, and to describe differences between the generalized cache and NC 10226. ATU 10228 and PC 10234 will then be described by description of differences between ATU 10228 and PC 10234 and the generalized cache.

Referring to Fig. 240, a partial block diagram of a generalized four-way, set associative cache is shown.

Tag Store (TS) 24010 is comprised of Tag Store A (TSA) S.24012, Tag Store B (TSB) 24014, Tag Store C (TSC) 24016, and Tag Store D (TSD) 24018. Each of the cache's sets, represented by TSA 24012 to TSD 24018, may contain, for example as in NC 10226, up to 16 entries, so that TSA 24012 to TSD 24018 are each 16 words long.

Address inputs to a cache are divided into a tag field and an index field. Tag fields are stored in the cache's tag store and indexed, that is addressed to be read or written from or to tag store by index field of the address. A tag read from tag store in response to

I

-273index field of an address is then compared to tag field of that address to indicate whether the cache contains an entry corresponding to that address, that is, whether a cache hit occurs. In, for example, NC 10226, a Name syllable may be comprised of an 8, 12, or 16 bit binary word, as previously described. The four least significant bits of these word or Names, comprise NC 10226's index field while the remaining,4, 8, or 12 most significant bits comprise NC 10226's tag field. TSA 24012 to TDS 24018 may each, therefore, be 12 entry wide memories to store the 12 bit tag fields of 16 bit names.

Index (IND) or address inputs of TSA 24012 to TSD 24018, would in NC 10226, be connected from four least t \significant bits of NAME Bus 20224 while Tag Inputs (TAGI) of TSA 24012 to TSD 24018 would be connected frcm the 12 most significant bits of NAME Bus 20224.

As described above, tag outputs of TS 24010 are compared to tag fields of addresses presented to the cache to determine whether the cache contains an entry 20 corresponding to that address. Using NC 10226 as an example 12 bit Tag Outputs (TAGOs) of TSA 24012 to TSD 24018 are connected to first inputs of Tag Store Comparators (TSC) 24019, respectively to inputs of Tag Store Comparitor A (TSCA) 24020, Tag Store Comparitor B (TSCB) 24022, Tag Store Comparitor D (TSCD) 24024, and Tag Store Comparitor E (TSCE) 24026. Second 12 bit inputs of TSCA 24020 to TSCE 24026 may be connected from the 12 most significant bits of NAME Bus 20224 to receive -274tag fields of NC 10226 addresses. TAS 24020 to TSCE 24026 compare tag field of an address to tag outputs read from TSA 24012 to TSE 24018 in response to index field of that address, and provide four bit outputs indicating which, if any, of the possible 16 entries and their associated tag store correspond to that address tag ofield. TSCA 24020 to TSCE 24026 may be comprised, for Sexample, of Fairchild 93S46s.

.r Four bit outputs of TSCA 24012 to TSCE 24026 are Sa 10 connected in the generalized cache to inputs of Tag Store SPipeline Registers (TSPR) 24027; respectively to inputs of Tag Store Pipeline Register A (TSPRA) 24028, Tag Store Pipeline Register B (TSPRB) 24030, Tag Store Pipeline Register C (TSPRC) 24032, and Tag Store Pipeline Register e C 15 D (TSPRD) 24034. ATU 10228 and PC 10234 is pipelined with a single cache access operation being executed in cc l two clock cycles. During first clock cycle tag store is addressed and tags store therein compared to tag field of address to provide indication of whether a cache hit has 20 occurred, that is whether cache contains an entry icorresponding to a particular address. During second clock cycle, as will be described below, a detected cache hit is encoded to obtain access to a corresponding entry in cache data store. Pipeline operation over two clock cycles is provided by cache pipeline registers which include, in part, TSPRA 24028 to TSPRD 24034. NC 10226 is not pipelined and does not include TSPRA 24028 to TSPRD 24034. In NC 10226, outputs of TSCA 24012 to TSCD -275- 24024 are connected directly to inputs of TSHEA 24036 to TSHED 24042, described below.

Outputs of TSPRA 24028 to TSPRD 24034 are connected to inputs of Tag Store Hit Encoders (TSHE) 24035, respectively to Tag Store Hit Encoder A (TSHEA) 24036, Tag Store Hit Encoder B (TSHEB) 24038, Tag Store Hit •l .Encoder C (TSHEC) 24040, and Tag Store Hit Encoder D (TSHED) 24042. TSHEA 24036 to TSHED 24042 encode, respectively, bit inputs from TSPRA 24028 to TSPRD 24034 to provide single bit outputs indicating which, if any, Sset of the cache's four sets includes an entry 1 corresponding to the address input.

Single bit outputs of TSHEA 24036 to TSHED 24042 are connected to inputs of Hit Encoder (HE) 24044. HE 24044 15 encodes single bit inputs from TSHEA 24036 to TSHED 24042 to provide two sets of ouputs. First outputs of HE 24044 S. are provided to Cache Usage Store (CUS) 24046 and indicate in which of the cache's four sets, corresponding to TSA 24012 to TSD 24018, a cache hit has occurred. As 20 described previously with reference to MC 20116, and will be described further below, CUS 24046 is a memory containing information for tracking usage of cache entries. That is, CUS 24046 contains entries indicating whether, for a particular index, Set A, Set B, Set C or Set D of the cache's four sets has been most recently used and which has been least recently used. CUS 24046 entries regarding Sets A, B, C, and D are stored in, respectively, memories CUSA 24088, CUSB 24090, CUSC

I-

L

U

-276- 24092, and CUSD 24094. Second output of HE 24044, as described further below, is connected to selection input of Data Store Selection Multiplexer (DSSMUX) 24048 to select an output from Data Store (DS) 24050 to be provided as output of the cache when a cache hit occurs.

Referring to DS 24050, as previously described a $o cache's data store contains the information, or entries, S^48 stored in that cache. For example, each entry in NC.

So 10226's DS 24050 is a logical descriptor comprised of an AON, and Offset, and Length. A cache's data store I parallels, in structure and organization, that cache's *tag store and entries therein are identified and located through that cache's tag store and associated tag store comparison and decoding logic. In NC 10226, for example, for each Name having an entry in NC 10226 there will be an entry, the tag field of that name, stored in TS 24010 and a corresponding entry, a logical descriptor corresponding to that Name, in DS 24050. As described above, NC 10226 is a four-way, set associative cache so I 20 that TS 24010 and DS 24050 will each contain four sets of tt, data. Each set was previously described as containing up to 16 entries. DS 24050 is therefore comprised of four 16 word memories. Each memory is 65 bits wide, accommodating 28 bits of AON, 32 bits of offset, and bits of length. These four component data store memories of DS 24050 are indicated in Fig. 240 as Data Store A (DSA) 24052, Data Store B (DSB) 24054, Data Store C (DSC) 24056, and Data Store D (DSD) 24058. DSA 24052, DSB -277- 24054, DSC 24056 and DSD 24058 correspond, respectively, in structure, contents, and operation to TSA 24012, TSB 24014, TSC 24016 and TSD 24018.

Data Inputs (DIs) of DSA 24052 to DSD 24058 are, in NC 10226 for example, connected from AON Bus 20230, OFFSET Bus 20228, LENGTH Bus 20226 and comprise inputs WA, WO, WL respectively of NC 10226. DSA 24052 to DSD h 24058 DIs are, in NC 10226 as previously described, utilized in writing NC 10226 entries into DSA 24052 to j 10 DSD 24058. Address inputs of DSA 24052 to DSD 24058 are connected from address outputs of Address Pipeline Register (ADRPR) 24060. As will be described momentarily, except during cache flush operations, DSA 24052 to DSD 24058 address inputs are comprised of the same index fields of cache addresses as are provided as address inputs to TS 24010, but are delayed by one clock cycle and ADRPR 24060 for pipelining purposes. As described above, NC 10226 is not pipelined and does not have the one clock cycle delay. An address input to the /20 cache will thereby result in corresponding entries, rtnlt selected by index field of that address, being read from TSA 24012 to TSD 24018 and DSA 24052 to DSD 24058. The four outputs of DSA 24052 to DSD 24058 selected by a particular index field of a particular address are provided as inputs to DSSMUX 24048. DSSMUX 24048 is concurrently provided with selection control input from HE 24044. As previously described, this selection input to DSSMUX 24048 is derived from TS 24010 t-a entries and -278indicates which of DSA 24052 to DSD 24058 entries corresponds to an address provided to the cache. In response to that selection control input, DSSMUX 24048 selects one of DS 24050's four logical descriptor outputs as the cache's output corresponding to that address.

DSSMUX 24048's output is then provided, through Buffer Driver (BD) 24062 as the cache's output, for example in NC 10226 to AON Bus 20230, OFFSET Bus 20228, and LENGTH Bus 20226.

Referring to ADRMUX 24062, ADRMUX 24062 selects one S of two sources to provide address inputs to DS 24050, that is to index to DS 24050. As described above, a first ADRMUX 24062 input is comprised of the cache's Saddress index fields and, for example in NC 10226, is connected from the four least significant bits of NAME Bus 20224. During cache flush operations, DS 24050 address inputs are provided from Flush Counter (FLUSHCTR) 24066, which in the example is a four bit counter.

During cache flush operations, FLUSHCTR 24066 generates Jo 20 sequential bit addresses which are used to sequentially address DSA 24052 to DSD 24058. Selection between ADRMUX 24062 first and second inputs, respectively the address index fields and from FLUSHCTR 24066, is controlled by Address Multiplexer Select (ADRMUXS) from FUCTL 20214.

Validity Store (VALS) 24068 and Dirty Store (DIRTYS) 24070 are memories operating in parallel with, and addressed in parallel wih TS 24010. VALS 24068 contains entries indicating validity of corresponding TS 24010 and DS 24050 entries. That is, VALS 24068 entries indicate whether corresponding entries have been written into 1 -279corresponding locations in TS 24010 and DS 24050. In the example, VALS 24068 may thereby be a 16 word by 4 bit wide memory. Each bit of a VALS 24068 word indicates validity of a corresponding location in TSA 24012 and DSA 24052, TSB 24014 and DSB 24054, TSC 24016 and DSC 24056, and TSD 24018 and DSD 24058. DIRTYS 24070 similarly indicates whether corresponding entries in corresponding locations of TS 24010 and DS 24050 havebeen written over, or modified. Again, DIRTYS 24070 will be a sixteen 10 word by four bit wide memory.

Address inputs of VALS 24068 and DIRTYS 24070 are, S. for example in NC 10226, connected from least significant bits of NAME Bus 20224 and are thus addressed by index fields of NC 10226 addresses in parallel with TS 24010.

Outputs of VALS 24068 are provided to TSCA 24020 to TSEE 24026 to inhibit outputs of TSCA 24020 through TSCE 24026 upon occurrence of an invalid concurrence between a TS 24010 entry and a NC 10226 address input. Similar *,F9 outputs of DIRTYS 24070 are provided to FUCTL 20214 for use in cache flush operations to indicate which NC 10226 entries are dirty and must be written back into an MT S' 10350 rather than disgarded.

Outputs of VALS 24068 and DIRTYS 24070 are also connected, respectively, to inputs of Validity Pipeline Register (VALPR) 24072 and Dirty Pipeline Register (DIRTYPR) 24074. VALPR 24072 and DIRTYPR 24074 are pipeline registers similar to TSPRA 24028 to TSPRD 24034 and are provided for timing purposes as will be described momentarily. Outputs of VALPR 24072 and DIRTYPR 24074 are connected to inputs of, respectively, Validity Write r; I -280tt.# t 4r &t i 4I1 4 I II tL I Logic (VWL) 24076 and Dirty Write Logic (DWL) 24078. As described above, NC 10226 is not a pipelined cache and does not include VALPR 24072 and DIRTYPR 24074; outputs of VALS 24068 and DIRTYS 24070 are connected directly to inputs of VWL 24076 and DWL 24078. Outputs of VWL 24076 and DWL 24078 are connected, respectively, to data inputs of VALS 24068 and DIRTYS 24070. Upon occurrence of a write operation to TS 24010 and DS 2405Q, that is writing in or modifying a cache entry, corresponding validity and dirty word entries are read from VALS 24068 and DIRTYS 24070 by index field of the caches input address.

Outputs to VALS 24068 DIRTYS 24070 are received and stored in, respectively, VALPR 24070 and DIRTYPR 24074.

At start of next clock cycle, validity and dirty words in VALPR 24072 and DIRTYPR 24074 are read into, respectively, VWL 24076 and DWL 24078. VWL 24076 and DWL 24078 respectively modify those validity or dirty word entries from VALS 24068 and DIRTYS 24070 in accordance to whether the corresponding entries in TS 24010 and DS 24050 are written into or modified. These modified validity and dirty words are then written, during second clock cycle, from VWL 24076 and DWL 24078 into, respectively, VALS 24068 and DIRTYS 24070. Control inputs of VWL 24076 and DWL 24078 are provided from FUCTL 25 20214.

Referring finally to Least Recent Used Logic (LRUL) 24080, LRUL 24080 tracks usage of cache entries. As previously described, the generalized cache of Fig. 240 1 1 u 1 -281is a four way, set associative cache with, for example, up to 16 entries in each of NC 10226's sets, Entries within a particular set are identified, as described above, by indexing the cache's TS 24010 and DS 24050 may contain, concurrently, up to four individual entries identified by the same index but distiguished by having different tags. In this case, one entry would reside in Set A, comprising TSA 24012 and DSA 24052, one in Set B, comprising TSB 24014 and DSB 24054, and so on. Since the of O' 10 possible number of individual entries having a common tag is greater than the number of cache sets, it may be necessary to delete a particular cache entry when another entry having the same tag is to be written into the cache. In general, the cache's least recently used entry would be deleted to provide a location in TS 24010 and DS 24050 for writing in the new entry. LRUL 24080 assists S" in determining which cache entries are to be deleted when necessary in writing in a new entry by tracking and indicating relative usage of the cache's entries. LRUL 24080 is primarily comprised of a mmory, LRU Memory (MLRU) 24081, containing a word for each cache set. As 4444 described above, NC 10226, for example, includes 16 sets of 4 frames each, so that LRUL 24080's memory may correspondingly be, for example, 16 words long. Each word indicates relative usage of the 4 frames in a set and is a 6 bit word.

Words are generated and written into LRUL 24080s MLRU 24081, through Input Register A, B, C, D (RABi^ 24083, according to a write only algorithm ex, !i 1 -282- 24044, as described momentarily. Each bit of each six word pertains to a pair of frames within a particular cache set and indicates which of those two frames was more recently used than the other. For example, Bit 0 will contain logic 1 if Frame A was used more recently than Frame B and a logic zero if Frame B was used more recently than Frame A. Similarly, Bit 1 pertains to Frames A and C, Bit 2 to Frames A and D, Bit 3 to Frames B and C, Bit 4 to Frames B and D, and Bit 5 to Frames C en" 10 and D. Initially, all bits of a particular LRUL 24080 te word are set to zero. Assuming, for example, that the frames of a particular set are used in the sequence Frame S t A, Frame D, Frame B- Bits 0 to 5 of that LRUL 24080 word will initially contain all zeros. Upon a reference to Fram A, Bits 0, 1, and 2, referring respectively to Frames A and B, Frames A and C, and Frames A and D, will t! t 1 be written as logic l's. Bits 3, 4, and 5, referring repectively to Frames B and C, Frames B and D, and Frames C and D, will remain logic 0. Upon reference to Frame D, Bits 0 and 1, referring respectively to Frames A and P and Frames A and C, will remain logic 1's. Bit 2, referring to Frames A and D, will be changed from logic 1 S to logic 0 to indicate that Frame D has been referred to S, more recently than Frame A. Bit 3, referring to Frames B and C, will remain logic 0. Bits 4 and 5, referring respectively to Frames B and D and Frames C and D, will be written as logic 0, although they are already logic zeros, to indicate respectively that Frame D has been used more recently than Frame B or Frame C. Upon reference to Frame B, Bit 0, referring to Frames A and B, -283will be written to logic 0 to indicate that Frame B has been used more recently than Frame A. Bits 1 and 2, referring resectively to Frames A and C and Frames A and D, will remain respectively as logic 1 and logic 0. Bits three and four, referring respectively to Frames B and C and Frames B and D, will be written as logics l's to indicate respectively that Frame B has been used more recently than Frame C or Frame D. Bit five will remain logic 0.

10 When it is necessary to replace a cache entry in a particular frame, the LRUL 24080 word referring to the ;cache set containing that frame will be read from LRUL S. 24080's MLRL 24081 through LRU Register (RLRU) 24085 and decoded by LRU Decode Logic (LRUD) 24087 to indicate which is least recently used frame. This decoding is executed by means of a Read Only Memory operating as a (set of decoding gating.

Having described the structure and operation of a S'.generalized cache as shown in Fig. 240, with references to NC 10226 for illustration and to point out differences between the generalized cache and NC 10226, structure and operation of ATU 10228 and PC 10234 will be described next below. ATU 10228 and PC 10234 will be described by describing the differences between ATU 10228 and PC 10234 and the generalized cache and NC 10226. ATU 10228 will be described first, followed by PC 10234.

-284d.d. Address--ransateion-it-rH- 2-28 and--Protection--eache-1- 4 ATU 10228 is a three-way, set associative cache of 16 sets, that is contains 3 frames for each set. Structure and operation of ATU 10228 is similar to the generalized cache described above. Having 3 rather than 4 frames per set, ATU 10228 does not include a STD 24018, ATSCE 24026, ATSPRD 24034, ATSHED 24042, or ADSD 24058. As previously described ATU 10228 address inputs comprise AON and 0 fields of logical descriptors. AON fields are each 28 bits and O fields comprise the 18 most significant bits of logical descriptor offset fields, so that ATU 10228 t' address inputs are 48 bits wide. Four least significant bits of 0 fields are used as index. AON fields and the 14 most significant bits of 0 field comprise ATU 10228's 1: tags. ATU 10228 tags are thereby each 42 bits in width.

S, Accordingly, TSA 24012{ TSB 24014, and TSC 24016 of ATU 10228's TS 24010 are each 16 words long by 42 bits wide.

S, DSA 24052, DSB 24054, and DSC 24056 of ATU 10228 are each 16 bits long. ATU 10228 outputs are, as previously described, physical descriptor Frame Number (FN) fields of 13 bits each. ATU 10228's DSA 24052, DSB 24054, DSC t 24056 are thereby each 13 bits wide.

ATU 10228's LRUL 24080 is similar in structure and operation to that of the generalized cache. ATU 12028's LRUL 24080 words, each corresponding to an ATU 10228 set, are each 3 bits in width as 3 bits are sufficient to indicate relative usage of frames within a 3 frame set.

*i i a~ 1^ -285- In ATU 10228, Bit 1 of an LRUL 24080 word indicates whether Frame A was used more recently than Frame B, Bit 2 whether Frame A was used more recently than Frame C, and Bit 3 whether Frame B was used more recently than Frame C. In all other respects, other than as stated above, ATU 10228 is similar in structure and operation to the generalized cache.

too Referring to PC 10234, PC 10234 is a two-way, set associative cache of 8 sets, that is has two frames per 10 set. Having 2 rather than 4 frames, PC 10234 will not Sinclude a TSC 24016, a TSD 24018, a TSCC 24024, a TSCD 24026, a TSPRC 24032, a TSPRD 24034, a TSHEC 24040, a TSHED 24042, a DSC 24056, or a DSD 24058.

Address inputs of PC 10234 are the 28 bit AON fields of logical descriptors. The 3 least significant bits of those AON fields are utilized as indexes for addressing PC 10234's TS 24010 and DS 24050. The 25 most significant.bits of those AON field address inputs are utilized as PC 10234's tags, so that PC 10234's TSA 24012 and TSB 24014 are each 8 word by 25 bit memories.

Referring to PC 10234's LRUL 24080, a single bit is sufficient to indicate which of the two frames in each of PC 10234's sets was most recently accessed. PC 10234's LRUL 24080's memory is thereby 8 words, or sets long, one bit wide.

As previously described, PC 10234 entries comprise o information regarding access rights of certain active subjects to certain active objects. Each PC 10234 entry contains 35 bits of information. Three bits of this information indicate whether a particular subject was -286read, write, or execute rights relative to a particular object. The remaining 32 bits effectively comprise a length field indicating the volume or portion,, that is the number of data bits, of that object to which those access rights pertain.

Referring again to Fig. 240, PC 10234 differs from the generalized cache and from NC 10226 and ATU 10228 in further including Extent Check Logic (EXTCHK) 24082 and Operation Check Logic (OPRCHK) 24084. PC 10234 entries 10 include, as described above, 3 bits identifying type of access rights a particular subject has to a particular object. These 3 bits, representing a Read Write St or Execute right, are provided to a first input of OPRCHK 24084. A second input of OPRCHK 24084 is provided from FUCTL 20214 and specifies whether JP 10114 intends to perform a Read a Write or Execute operation with respect to that object. OPRCHK 24084 compares OPRCHK 24084 access right inputs from DS 24050 to OPRCHK 24084's intended operation input from FUCTL 20214. If that subject does not possess the rights to that object which are required to perform the operation intended by JP 10114, OPRCHK 24084 generates an Operation Violation (OPRV) indicating that a protection violation has occurred.

Similarly, the 32 bits of a PC 10234 entry regarding extent rights is provided as an input (EXTENT) to EXTCHK 24082. As stated above, EXTENT field of PC 10234 entry indicates the length or number of data bits, within an obect, to which those access rights pertain. EXTENT field from PC 10234 entry is compared, by EXTCHK 24082,

L

-287to offset field of the logical descriptor of the current JP 10114 request to MEM 10112 for which a current protection mechanism check is being made. If comparison of extent rights and offset field indicate that the current memory request goes beyond the object length to which the corresponding rights read from DS 24050 apply, EXTCHK 24082 generates an Extent Violation (EXTV) output. EXTV indicates that a current memory request by JP 10114 refers to a portion of an object to which the PC 10234 entry read from BS 24050 does not apply. As described previously, each read from or write to MEM 10112, even as part of a string transfer, is a 32 bit word. As such, EXTCHK 24082 will generate an EXTV output Owhen OFFSET field of a current logical descriptor describes a segment of an object less than 32 bits from the limit defined by EXTENT field of the PC 10234 entry &a provided in response to that logical descriptor. EXTV and OPRV are gated together, by Protection Violation Gate 'a .(PVG) 24086 to generate Protection Violation (PROTV) output indicating that either an extent or an operation violation has occurred.

Having described the structure and operation of MEMINT 20212, and previously the structure and operation of DESP 20210, structure and operation of FUCTL 20214 will be described next below.

Si fc° -288- 3. Fet c h- -Un-i--eon4t~o--oqe--2-2+4--+Rg-.

The following descriptions will provide a detailed description of FU 10120's structure and operation.

Overall operation of FU 10120 will be described first, followed by description of FU 10120's structure, and finally by a detailed description of FU 10120 operation.

As previously described, FUCTL 20214 directs operation of JP 10114 in executing procedures of user's processes. Among the functions performed by FUCTL 20214 are, first, maintenance and operation of CS 10110's Name Space, UID, and AON based addressing system, previously *described; second, interpretation of SOPs of user's processes to provide corresponding sequences of microinstructions to FU 10120 and EU 10122 to control operation of JP 10114 in execution of user's processes, S*.o previously described; and, third, control of operation of CS 10110's internal mechanisms, for example CS 10110's stack mechanisms.

As will be described in further detail below, FUCTL 20214 includes Prefetcher (PREF) 20260 which generates a sequence of logical addresses, each logical address 0 comprising an AON and an offset field, for reading S- Instructions (SINs) of a user's program from MEM 10112.

As previously described, each SIN may be comprised of an S-Operation (SOP) and one or more operand Names and may occupy one or more 32 bit words. SINs are read from MEM 10112 as a sequence of single 32 bit words, so that PREF -289- 20260 need not specify a length field in a MEM 10112 read request for an SIN. SINs are read from MEM 10112 through MOD Bus 10144 and are captured and stored in Instruction Buffer (INSTB) 20262. PARSER 20264 extracts, or parses, SOPs and operand Names from INSTB 20262. PARSER 20264 provides operand Names toNC 10226 and SOPs to FUS Intrepreter Dispatch Table (FUSDT) 11010 and to EU Dispatch Table (EUSDT) 20266 through Op-Code Register (OPCODEREG) 20268. Operation of INSTB 20262 and PARSER f* 10 20264 is controlled by Current Program Counter (CPC) 20270, Initial Program Counter (IPC) 20272, and Executed Program Counter (EPC) 20274.

As previously described, FUSDT 11010 provides, for "each SOP received from OPCODEREG 20268, a corresponding S-Interpreter Dispatch (SD) Pointer, or address, to FUSITT 11012 to select a corresponding sequence of O' microinstructions to direct operation of JP 10114, in 4 particular FU 10120. As previously described, FUSITT 11012 also contains sequences of microinstructions for 20 controlling and directing operation of CS 10110's internal mechanisms, for example those mechanisms such as RCWS 10358 which are involved in swapping of processes.

EUSDT 20266 performs an analogous function with respect t to EU 10122 and provides SD Pointers to EU S-Interpreter I Tables (EUSITTs) residing in EU 10122.

Micro-Program Counter (mPC) 20276 provides sequential addresses to FUSITT 11012 to select individual microinstructions of sequences of microinstructions.

-290- Branch and Case Logic (BRCASE) 20278 provides addresses to FUSITT 11012 to select microinstructions sequences for microinstructions branches and and cases. Repeat Counter (REPCTR) 20280 and Page Number Register (PNREG) 20282 provide addresses to FUSITT 11012 during FUSITT 11012 load operations.

As previously described, FUSITT 11012 is a writable microinstruction control store which is loaded with selected S-Interpreters (SINTs) from MEM 10112.

FUSITT 11012 addresses are also provided by Event Logic (EVENT) 20284 and by JAM input from NC 10226. As Swill be described further below, EVENT 20284 is part of FUCTL 20214's circuitry primarily concerned with St operation of CS 10110's internal mechanisms. Input JAM from NC 10226 initiates certain FUCTL 20214 control functions for CS 10110's Name Space addressing td; mechanisms, and in particular NC 10226. Selection between the above discussed address inputs to FUSITT 11012 is controlled by S-Interpreter Table Next Address 0 20 Generator Logic (SITTNAG) 20286.

Other portions of FUCTL 20214's circuitry are concerned with operation of CS 10110's internal S mechanisms. For example, FUCTL 20214 includes Return Control Word Stack (RCWS) 10358, previously described with reference to CS 10110's Stack Mechanisms. Register Address Generator (RAG) 20288 provides pointers for addressing of GRF 10354 and RCWS 10358 and includes Micro-Stack Pointer Registers (MISPR) 10356.

-291- As previously described, MISPR 10356 mechanism provides pointers for addressing Micro-Stack (MIS) 10368. As will be described further below, actual MIS 10368 Pointers pointing to current, previous, and bottom frames of MIS 10368 reside in Micro-Control Word Register 1 (MCW1) 20290. MCW1 20290 and Micro-Control Word Zero Register (MCWO 20292 together contain certain information indicating the current execution environment of a microinstruction sequence currently being executed o**a 10 by FU 10120. This execution information is used in aide of execution of these microinstruction sequences. State Registers (STATE) 20294 capture and store certain information regarding state of operation of FU 10120. As S, described further below, this information, referred to as 15 state vectors, is used to enable and direct operation of FU 10120.

Timers (TIMERS) 20296-monitor elapsed time since occurrence of the events requiring servicing by FU 10120.

If waiting time for these events exceeds certain limits, 20 TIMERS 20296 indicate that these limits have been exceeded so that service of those events may be initiated.

Finally, Fetch Unit to E Unit Interface Logic c t i (FUEUINT) 20298 comprises the FU 10120 portion of the interface between FU 10120 an EU 10122. FUEUINT 20298 is primary path through which operation of FU 10120 and EU 10122 is coordinated.

Having described overall operation of FU 10120, structure of FU 10120 will be described next below with

L

r r, ii I~_L CI-L- llC14 i 11~1_3~ L -292-

S

cc I Sc

II

S I Srf S I' #5 St i4** *i 4* 5 5 *c 5 5

S

*c 4 *c S aide of Fig. 202, description of FU 10120's structure will be followed by a detailed description of FU 10120 wherein further, more detailed, diagrams of certain portions of FU 10120 will be introduced as required to enhance clarity of presentation.

a.a. Petch--Un+t--eontro--L- e--2214 Overa-l--St-ucttue Referring again to Fig. 202, as previously described Fig. 202 includes a partial block diagram of FUCTL 10 20214. Following the same sequence of description as above, PREF 20260 has a 28 bit bi-directional port connected to AON Bus 20230 and 32 bit bi-directional port directed from OFFSET Bus 20228. A control input of PREF 20260 is connected from control output of INSTB 20262.

15 INSTB 20262 32 bit data input (DI) is connected from MOD Bus 10144. INSTB 20262's 16 bit output (DO) is connected to 16 bit bi-directional input of OPCODEREG 20268 and to 16 bit NAME Bus 20224. OPCODEREG 20268's input comprises 8 bits of SINT and 3 bits of dialect 20 selection. As previously described, NAME Bus 20224 is connected to 16 bit bi-directional port of Name Trap (NT) 20254, to address input ADR of NC 10226, and to inputs and outputs of OFFP 20228. Control inputs of INSTB 20262 and PARSER 20264 are connected from a control output of CPC 20270.

Thirty-two bit input of CPC 20270 is connected from JPD Bus 10142 and CPC 20270's 32 bit output is connected to 32 bit input of IPC 20272. Thirty-two bit output of IPC 20272 is connected to 32 bit input of EPC 20274 and ElJ1 i^f i -293to JPD Bus 10142. EPC 20274's 32 bit output is similarly connected to JPD Bus 10142.

Eleven bit outputs of OPCODEREG 20268 are connected to 11 bit address inputs of FUSDT 11010 and EUSDT 20266.

These 11 bit address inputs to FUSDT 11010 and EUSDT 20266 each comprise 3 bits of dialect selection code and 8 bits of SINT code. Twelve bit SDT outputs of EUSDT 20266 is connected to inputs of Microinstruction Control Store in EU 10122, as will be described in a following I T" 10 description of EU 10122. FUSDT 11010 has, as described S* further below, two outputs connected to address (ADR) Bus 20298. First output of FUSDT 11010 are six bit SDT pointers, or addresses, corresponding to generic SINTs as will be described further below. Second output of FUSDT 11010 are 15 bit SD^ pointers, or addresses, for algorithm microinstruction sequences, again as will be described further below.

«..Referring to RCWS 10358, RCWS 10358 has a first bidirectional port connected from JPD Bus 10142. Second, 30 third, and fourth bi-directional ports of RCWS 10358 are connected from, respectively, a bi-directional port of MCW1 20290, a first bi-directional port EVENT 20284, and a bi-directional port of mPC 20276. An output of RCWS 10358 is connected to ADR Bus 20298.

An input of mPC 20276 is connected from ADR Bus 20298 and first and second outputs of mPC 20276 are connected to, respectively, an input of BRCASE 20278 and to ADR Bus 20298. An output of BRCASE 20278 is connected to ADR Bus 20298.

'u -294- As described above, a first bi-directional port of EVENT 20284 is connected to RCWS 10353. A second bidirectional port of EVENT 20284 is connected from MCWO 20292. An output of EVENT 20284 is connected to ADR Bus 20298.

Inputs of RPCTR 20280 and PNREG 20282 are connected from JPD Bus 10142. Outputs of RPCTR 20280 and PNREG 20282 are connected to ADR Bus 20298.

ADR Bus 20298, and an input from a first output of 10 FUSITT 11012, are connected to inputs of SITTNAG 20286.

.Output of SITTNAG 20286 is connected, through Control Store Address (CSADR) Bus 20299, to address input of FUSITT 11012. Data input of FUSITT 11012 is connected from JPD Bus 10142. Control outputs of FUSITT 11012 are connected to almost all elements of JP 10114 and thus, for clarity of presentation, are not shown in detail by drawn physical connections but are described in following descriptions.

As described above, MCWO 20292 and MCW1 20290 have S' 20 bi-directional ports connected to, respectively, bidirectional ports of EVENT 20284 and to a second bidirectional port of RCWS 10358. Outputs of MCWO 20292 and MCW1 20290 are connected to JPD Bus 10142. Other i inputs of MCWO 20292 and MCW1 20290, as will be described further below, are connected from several other elements of JP 10114 and, for clarity of presentation, are not shown herein in detail but are described in the following i text. STATE 20294 similarly has a large number of inputs

L

-295and outputs connected from and to other elements of JP 10114, and in particular FU 10120. Inputs and outputs of STATE 20294 are not indicated here for clarity of presentation and will be described in detail below.

RAG 20288 has an input connected from JPD Bus 10142 and other inputs connected, for example, from MCW1 20290. RAG 20288, including MISPR 10356, provides outputs, for example, as address inputs, to RCWS 10358 and GRF 10354. Again, for clarity of presentation, inputs 10 and outputs of RAG 20288 are not shown in detail in Fig.

202 but will be described in detail further below.

TIMERS 20296 receive inputs from EVENT 20284 and FUSITT 11012 and provide outputs to EVENT 20284. For clarity of presentation, these indications are not shown in detail in Fig. 202 but will be described further below.

FUINT 20298 receives control inputs from FUSITT 11012 and EU 10122. FUINT 20298 provides outputs to EU 10122 and to other elements of FUCTL 20214. For clarity of 20 presentation, connections to and from FUINT 20298 are not shown in detail in Fig. 202 but will be described in further detail below.

Having described the overall operation, and structure, of FUCTL 20214, operation of FUCTL 20214 will be described next below. During the following descriptions further diagrams of certain portions of FUCTL 20214 will be introduced as required to disclose structure and operation of PUCTL 20214 to one of ordinary -296skill in the art. FUCTL 20214's operation with regard to fetching and interpretation of SINs, that is SOPs and operand Names, will be described first, followed by description of FUCTL 20214's operation with regard to CS 10110's internal mechanisms.

b.b. Petch- Unit--eontro-l--oa+T--20-2-4 Ope-faon Referring first to those elements of FUCTL 20714 directly concerned with dontrol of JP 10114 in response o10 to SOPs and Name syllables, those elements include: (1) PREF 20260; INSTB 20262; PARSER 20264; CPC 20270, IPC 20272, and EPC 20274; OPCODEREG 20268; (6) FUSDT 11010 and EUSDT 20266; mPC 20276; BRCASE 20278; REPCTR 20280 and PNREG 20282; (10) a part of RCWS 10358; (11) SITTNAG 20286; (12) FUSITT 11012; and, (13) NT 20254. These FUCTL 20214 elements will be S^ described below in the ordf=r named.

Sa.aa.a. Pref etch er--20-6t Re2fei-20-2G P -RPt-2-22-2-and8-e--2G-2-7-4 4r gFfizP--24i- 4 ta: As described above, PREF 20260 generates a series of addresses to MEM 10112 to read SINs of user's programs from MEM 10112 to FUCTL 20214, and in particular to INSTB 20262. Each PREF 20260 read request transfers one 32 bit word from MEM 10112. Each SIN may be comprised of an SOP and one or more Name syllables. Each SOP may comprise, for example, 8 bits of information while each Name -297syllable may comprise, for example, 8, 12, or 16 bits of data. In general, and as will be described in further detail in a following description of STATE 20294, PREF 20260 obtains access to MEM 10112 on alternate 110 nanosecond system clock cycles. PREF 20260's access to MEM 10112 is conditional upon INSTB 20262 indicating that INSTB 20262 is ready to receive an SIN read fsom MEM 10112. In particular, INSTB 20262 generates control output Quiry Prefetch (QPF) to PREP 20260 to enable PREF 10 20260 to submit a request to MEM 10112 when, as described 9. .further below, INSTB 20262 is ready to receive an SIN read from MEM 10112.

PREF 20260 is a counter register comprised, for example of SN74S163s.

r' 15 Bi-directional inputs and outputs of PREF 20260 are connected to AON Bus 20230 and OFFSET Bus 20228. As PREF 20260 reads only single 32 bit words, PREF 20260 is not 4 4* trequired to specify a LENGTH field as part of an SIN read request, that is an AON and an OFFSET field are sufficient to define a single 32 bit word. At start of S" read of a sequence of SINs from MEM 10112, address (AON and OFFSET fields) of first 32 bit word of that SIN sequence are provided to MEM 10112 by DESP 20210 and Sconcurrently loaded, from AON Bus 20230 and OFFSET Bus 20228, into PREF 20260. Thereafter, as each successive thirty-two bit word of the SIN's sequence is read from MEM 10112, the address residing in PREF 20260 is in(remented to specify successive 32 bit words of that

SI

1 I ir- 8 f yifp .K t -298- SIN'S sequence. The successive single word addresses are, for all words after first word of a sequence, provided to MEM 10112 from PREF 20260.

As described above, INSTB 20262 receives SINs from MEM 10112 through MOD Bus 10144 and, with PARSER 20264 and operating under control of CPC 20270, provides Name syllables to NAME Bus 20224 and SINs to OPCODEREG 20268.

INSTB 20262 is provided, together with PREF 20260 to increase execution speed of SINS.

0 Referring to Fig. 241, a more detailed block diagram S.s of INSTB 20262, PARSER 20264, CPC 20270, IPC 20272, EPC 20274 as shown. INSTB 20262 is shown as comprising two 32 bit registers having parallel 32 bit inputs from MOD S' Bus 10144. INSTB 20262 also receives two Write Clock (WC) inputs, one for each 32 bit register of INSTB 20262, from Instruction Buffer Write Control (INSTBWC) 24110.

SC INSTB 20262's outputs are structured as eight, eight bit 4. Basic Syllables (BSs), indicated as BSO to BS7. BSO, BS2, BS4, and BS6 are ORed to comprise eight bit Basic A' 20 Syllable, Even (BSE) of INSTB 20262 while BSO, BS3, and BS7 are similarly ORed to comprise Basic Syllable, Odd (BSO) of INSTB 20262. BSO and BSE are provided as I inputs of PARSER 20264.

PARSER 20264 receives a first control input from Current Syllable Size Register (CSSR) 24112, associated with CPC 20270. A second control input of PARSER 20264 is provided from Instruction Buffer Syllable Decode Register (IBSDECR) 24114, also associated with CPC 20270. PARSER 20264 provides an eight bit output to NAME Bus 20224 and to input of OPCODEREG 20268.

iI -299- Referring to INSTBWC 24110, INSTBWC 24110 provides, as described further below, control signals pertaining to writing of SINs into INSTB 20262 from MOD Bus 10144.

INSTBWC 24110 also provides control signals pertaining to operation of PREF 20260. In addition to WC outputs to INSTB 20262, INSTBWC 24110 provides control output QPF to PREF 20260, control output Instruction Buffer Hung (IBHUNG) to EVENT 20284, and control signal Instruction Buffer Wait (IBWAIT) to STATE 20294. INSTBWC 24110 also ol.'s 10 receives a control input BRANCH from BRCASE 20278 and an error input from TIMERS 20296.

t Referring to CPC 20270, IPC 20272, and EPC 20274, IPC t20272 and EPC 20274 are represented in Fig. 241 as in I r L 4 Fig. 202. Further FUCTL 20214 circuitry is shown as associated with CPC 20270. CPC 20270 is a twenty-nine bit register receiving bits one to twety-five (CPC(1-25)) SS from bits one to twenty-five of JPD Bus 10142. CPC 20270 Bit 0 (CPCO) is provided from CPCO CPCO Select (CPCOS) 24116. Inputs of CPCOS 24116 are Bit 1 output from CPC S, 20 20270 (CPC1) and Bit 0 from JPD Bus 10142. Bits twentysix, twenty-seven, and twenty-eight of CPC 20270 (CPC(26- 28)) are provided from CPC Multiplexer (CPCMUX) 24118.

I .I S CPCMUX 24118 also provides an input to IBSDECR 24114.

Inputs of CPCMUX 24118 are bits twenty-five, twenty-six, and twenty-eight from JPD Bus 10142 and a three bit output of CPC Arithmetic and Logic Unit (CPCALU) 24120.

A first input of CPCALU 24120 is connected from output bits 26, 27, and 28 of CPC 20270. Second input of CPCALU k: r 1'I -300- *4 4 4.

C: 4 24120 is connected from CSSR 24112. CSSR 24112's input is connected from JPD Bus 10142.

As described above, INSTB 20262 is implemented as a sixty-four bit wide register. INSTB 20262 is organized as two thirty-two bit words, referred to as Instruction Buffer Word 0 (IBO) and Instruction Buffer Word 1 (IBl), and operates as a two word, first-in-first-out buffer memory. PREF 20260 loads one of IBO or IB1 on each memory reference by PREF 20260. Only PREF 20260 may load 10 INSTB 20262, and INSTB 20262 may be loaded only from MOD Bus 10144. Separate clocks, respectively Instruction Buffer Write Clock 0 (IBWCO) and Instruction Buffer Write Clock 1 (IBWC1), are provided from INSTBWC 24110 to load, respectively, IBWO and IBW1 into INSTB 20262. IBWCO and IBWC1 are each a gated 110 nano-second clock. An IBWO or an IBW1 is written into INSTB 20262 when, respectively, IBWCO or IBWC1 is enabled by INSTBWC 24110. IBWCO and IBWCl will be enabled only when MEM 10112 indicates that data for INSTB 20262 is availabe by asserting interface control signal DAVI as previously discussed.

INSTBWC 24110 is primarily concerned with control of FU 10120 with respect to writing of SINs into INSTB 20262. As described above, INSTBWC 24110 provides IBWCO and IBWC1 to INSTB 20262. IBWCO and IBWC1 are enabled by INSTBWC 24110's input DAVI from MEM 10112. Selection between IBWCO and IBWC1 is controlled by INSTBWC 24110's input from CPC 20270. In particular, and as will be described further below, Bit 26 (CPC 26) of CPC 20270's twenty-nine bit word indicates whether IBWO or IBW1 is written into INSTB 20262.

r

:I

\gi ~1 if; f IK f* -301- In addition to controlling writing of IBWO and IBWl into INSTB 20262, INSTBWC 24110 provides control signals to elements of FU 10120 to control reading of SINs from MEM 10112 to INSTB 20262. In this regard, INSTBWC 24110 detects certain conditions regarding status of SIN words in INSTB 20262 and provides corresponding control signals, described momentarily, to other elements of FU 10120 so that INSTB 20262 would generally always contain at least one valid SOP or Name syllable. First, if INSTB 20262 is not full, that is either IBWO or IBWl or both is invalid, for example because IBWO has been read from INSTB 20262 and executed, INSTBWC 24110 detects this condition and provides control signal QPF to PREF 20262 to initiate a read from MEM 10112. INSTBWC 24110 currently enables either IBWO or IBW1 portion of INSTB 20262 to receive the word read from MEM 10112 in response to PREF 20260's request. As stated above, this operation S'will be initiated when INSTBWC 24110 detects and indicates, by generating a validity flag, that either IBWO or IBW1 is invalid. In this case, IBWO or IBWl will be indicated as invalid when read from INSTB 20262 by PARSER 20264. As will be described further below, INSTBWC 24110 validity flags for IBWO and IBWl ar generated by INSTBWC 24110 control inputs comprising Bits 26 to 28 (CPC 26-28) from CPC 20270 and by current 2A syllable size or value, flag input from CSSR 24112.

Secondly, INSTBWC 24110 will detect when INSTB 20262 is empty, that is when both IBWO and IBWl are invalid, as

IK

j s««:t L t beidctda nvldwe edfrmISB222b -302ft** 94 4 4 04* •o 4 4 4 9 4 4,* 9 *4* just described, or when only a half of a sixteen bit Name syllable is present in INSTB 20262. In response to either condition, INSTBWC 24110 will generate control signal IBWAIT to STATE 20294. As will be described further below, IBWAIT will result in suspension of execution of microinstructions referencing INSTB 20262.

PREF 20260 requests to MEM 10112 will already have been initiated, as described above unless certain other, conditions, described momentarily, occur. Thirdly, 10 INSTBWC 24110 will detect when INSTB 20262 is empty and PREF 20262 is hung, that is unable to submit requests to MEM 10112, and a current microinstruction is attempting to parse a syllable from INSTB 20262. In this case, INSTBWC 24110 will generate control signal Instruction 15 Buffer Hung (IBHUNG) to EVENT 20284. As will be described further below, IBHUNG will result in initiation of a microinstruction sequence to restore flow of words to INSTB 20262. Fourthly, INSTBWC 24110 will detect, through microinstruction control signals provided from FUSITT 11012, when a branch in a microinstruction sequence provided by FUSITT 11012 in response to an SOP occurs. In this case, both IBWO and IBWl will be flaged as invalid. INSTBWC 24110 will then ignore SIN words being read from MEM 10112 in response to a previously submitted PREF 20260 request, but not yet received at the time the branch occurs. This prevents INSTB 20260 from receiving invalid SIN words; PREF 20260 and INSTB 20262 will then proceed to request and receive valid SIN words of the branch.

-303- As described above, PARSER 20264, operating under control of CPC 20270 and CPC 20270 associated circuitry, reads Name syllables and SOPs from INSTB 20262 to, respectively, NAME Bus 20224 and OPCODEREG 20268. PARSER 20264 operates as a multiplexer with associated control logic.

As previously described, INSTB 20262 is internally structured as eight, eight bit words, BSO to BS7. IBWO comprises BSO to B3 while IBWI comprises BS4 to BS7.

I: 10 Each SOP is comprised of eight bits of data and thus comprises one Basic Syllable while each Name syllable S, comprises 8, 12, or 16 bits of data and thus comprises either one or two Basic Syllables. Name syllable size, as previously stated, is indicated by Current Syllable S 15 Size Value K stored in CSSR 24112.

I °BSO and BS4 are loaded into INSTB 20262 from MOD Bus S10144 bits zero to seven while BS2 and BS6 are loaded from MOD Bus 10144 bits sixteen to twenty-three. BS1 and are loaded from MOD Bus 10144 bits eight to fifteen 20 while BS3 and BS7 are loaded from MOD Bus 10144 bits *e twenty-four to thirty-one. Odd numbered Basic Syllable S* i outputs BS1, BS3, BS5, and BS7 are ORed to comprise eight bit Basic Syllable, Odd output BSO of INSTB 20262. Even numbered Basic Syllable outputs BSo, BS2, BS4 and BS6 of INSTB 20262 are similarly ORed to comprise eight bit Basic Syllable, Even output BSE. At any time, one odd numbered Basic Syllable output and one even numbered Basic Syllable output of INSTB 20262 are selected as inputs to PARSER 20264 by Instruction Buffer Read Enable (IBORE) enable and selection signals provided to INSTB IL".w i i f -304- 20262 by IBSDECR 24114. IBSDECR 24114 includes decoding circuitry. Input to IBSDECR 24114's decoding logic is comprised of three bits (RCPC(26-28)) provided from CPCMUX 24118. As indicated in Fig. 241, CPC (26-28) may be provided from JPD Bus 10142 bits 25 to 28 or from output of CPCALU 24120. One input CPCALU 24120 is CPC (26-28) from CPC 20270. Operation of CPC 20270 and CPC 20270's associated circuitry will be described further below. RCPC (26-28) is decoded by IBSDECR 24114 to 10 generate IBORE to INSTB 20262. RCPC 26 and RCPC 27 "are decoded to select one of the four odd numbered Basic Syllable outputs (that is BS1, BS3, BS5 or BS7) of INSTB 20262 as the odd numbered basic syllable input to PARSER 20264. RCPC 28 selects either the preceding or the tvf" 15 following even numbered Basic Syllable output of INSTB t C 20262 as the even numbered Basic Syllable input to PARSER 20264. The eight decoded bits of IBORE generated t by IBDECR 24114 decoding logic are loaded into IBSDECR 24114 eight bit register and subsequently provided to INSTB 20262 as IBORE SPARSER 20264 selects BSO, or BSE, or' both BSO and

S

t BSE, as PARSER 20264s output to NAME Bus 20224 or to OPCODEREG 20268. In the case of an SOP or an eight bit Name syllable, either BSO or BSE will be selected as PARSER 20264's output. In the case of a twelve or sixteen bit Name syllable, both BSO and BSE may be selected as PARSER 20264's output. PARSER 20264 operation is controlled by microinstruction control outputs from FUSITT 11012.

-305- Program counters IPC 20272, EPC 20274, and CPC 20270 are associated with control of fetching and parsing of SINs. In general, IPC 20272, EPC 20274, and CPC 20270 operate under microinstruction control from FUSITT 11012.

CPC 20270 is Current Program Counter and contains 28 bits pointing to the current syllable in INSTB 20272.

Bits 29 to 31 of CPC 20270 are not provided, so the bits 29 to 31 of CPC 20270's output are zero, which guarantees byte boundaries for SOPs. Contents of CPC 20270 are 10 thereby also a pointer which is a byte align offset into •a current procedure object. Initial Program Counter (IPC) 20272 is a buffer register connected from output of CPC 20270 and provided for timing overlap. IPC 20272 may be loaded only from CPC 20270 which, as previously 15 described, is 29 bits wide, that does not contain bits 29, 30, and 31 which are forced to zero in IPC 20272.

S IPC 20272 may be read onto JPD Bus 10142 as a start value I in an unconditional branch.

EPC 20274 is a thirty-two bit register usually S 20 containing a pointer to the current SOP being executed.

Upon occurrence of an SOP branch, the pointer in EPC I 9 20274 will point to the SOP from which the branch was executed. The pointer residing in EPC 20274 is an offset into a current procedure object. EPC 20274 may be loaded only from IPC 20272, and may be read onto JPD Bus 10142.

Referring again to CPC 20270, as described above CPC 20270 is a current syllable counter. CPC 20270 contains a pointer to the next SOP syllable, or Base Syllable, to be parsed by PARSER 20264. As SOPs are always on byte boundaries, CPC 20270 pointer is 29 bits wide, CPC (0- 1 i pr -306- 28). The three low order bits of CPC 20270's pointer, that is CPC (29-31), do not physically exist and are assumed to be always zero. CPC 20270's pointer to next instruction syllable to be parsed thereby always points to byte boundaries.

CPC 20270 bits 26 to 28, CPC (26-28), indicate, as described above, a particular Base Syllable in INSTB S20262. Bits 0-25 (CPC(0-25)) of CPC 20270 indicate 32 bit words, read into INSTB 20262 as IBWO and IBW1, of a So* 10 sequence of SINs. CPC 20270 pointer is updated each time t a parse operation reading a Base Syllable from INSTB 20262 is executed. As previously described, these parsing operations are performed under microinstruction control from FUSITT 11012.

15 Conceptually, CPC 20270 is organized as a twenty-six t bit counter, containing CPC with a three bit register appended on the low order side, as CPC (26-28).

'This organization is used because CPC (26-28) counts INSTB 20262 Base Syllables parsed and must be incremented dependant upon current Name Syllable Size K stored CSSR 24112. CPC however, counts successive thirty-two S*bit words of a sequence of SINs and may thereby be Simplemented as a binary counter. As shown in Fig. 241, CPC (26-28) is loaded from output of CPCMUX 24118. A t first input of CPCMUX 24118 is connected from bits 29 to 31 of JPD Bus 10142. This input to CPC (26-28) from JPD Bus 10142 is provided to allow CPC 20270 to be loaded from JPD Bus 10142, for example when loading CPC 20270 -307with an initial ponter value. Second input of CPCMUX 24118 is from output of CPCALU 24120 and is the path by which CPC (26-28) is incremented as successive Base Syllables are parsed from INSTB 20262. A first 'input of CPCALU 24120 is CPC (26-28) from CPC 20270. Second input of CPALU 24120 is a dual input from CSSR 24112. First input from CSSR 24112 is logic 1 in the least significant bit position, that is in position corresponding to CPC This input is used when single Base Syllables are parsed from INSTB 20262, for example in an eight bit SOP Sor an eight bit Name syllable. CSSR 24112's first input to CPCALU 24120 increments CPC (0-32) by eight, that is ,,cc one to CPC (26-28), each time a single Base Syllable is parsed from INSTB 20262. Second input to CPCALU 24120 from CSSR 24112 is K, that is current Name Syllable l size. As previously described, K may be eight, twelve, or sixteen. CPC (26-28) is thereby incremented by one when K equals eight and is incremented by two when K equals twelve or sixteen. As shown in Fig. 241, K is loaded into CSSR 24112 from JPD Bus 10142.

CPC as described above, operates as a twentysix -bit counter which is incremented each time CPC (26- S' 28) overflows. CPC (0-25) is incremented by carry output of CPCALU 24120. In actual implementation, CPC 20270 is organized to reduce the number of integrated circuits required. CPC (1-25) is constructed aa a counter and inputs of CPC (1-25) counter are connected from bits 1 to 24 of JPD Bus 10142 to allow loading of an initial value h

I

-308of CPC 20270 pointer. CPC and CPC (26-28) are implemented as a four bit register. Operation of CPC (26-28) portions of this register have been described above. Input of CPC portion of this register is connected from output of CPCOS 24116. CPCOS 24116 is a multiplexer having a first input connected from bit 0 of JPD Bus 10142. This input from JPD Bus 10142 is used, for example, when loading CPC 20272 with an initial pointer value. Second input of CPCOS 24116 is overflow 10 output of CPC (1-25) counter and allows CPC portion of the four bit register and CPC (1-25) counter to 1 ,0i operate as a twenty-six bit counter.

'r Finally, as shown in Fig 241, output of CPC 20270 may be loaded into IPC 20272. An initial CPC 20270 pointer value may therefore be written into CPC 20270 from JPD j Bus 10142 and subsequently copied into IPC 20272.

Referring again to PARSER 20264, as described above PARSER 20264 reads, or parses, basic syllables from INSTB S* 20262 to NAME Bus 20224. Input of PARSER 20264 is a sixteen bit word comprised of an eight bit odd numbered "Base Syllable, BSO, and an eight bit even numbered Base Syllable, BSE. Depending upon whether PARSER 20264 is .parsing an eight bit SOP, an eight bit Name syllable, a twelve bit Name syllable, or sixteen bit Name syllable, PARSER 20264 may select BSO, BSE, or both BSO and BSE, as output onto NAME Bus 20224. A Akibi s -309- If PARSER 20264 is parsing Name syllables and K is not equal to eight, that is equal to twelve or sixteen, PARSER 20264 transfers both BSO and BSE onto NAME Bus 20224 and determines which of BSO or BSE is most significant. The decision as to whether BSO or BSE is most significant is determined by CPC If CPC (28) indicates BSO is most significant, BSO is transferred onto NAME Bus 20224 bits 0 to 7 (NAME(0-7)) and BSE onto NAME Bus 20224 bits eight to fifteen (NAME(8-15)). If S 10 CPC (28) indicates BSE is most significant, BSE is transferred onto NAME and BSO onto NAME (8-15).

r 1 This operation insures that Name syllables are parsed onto NAME Bus 20224 in the order in which occur in the SIN stream.

If PARSER 20264 is parsing Name syllables of Syllable Size K 8, PARSER 20264 will select either BSO or BSE, S* as indicated by CPC as output to NAME PARSER 20264 will place O's on NAME (8-15).

4 a If PARSER 20264 is parsing SOPs of eight bits, PARSER 20264 will select BSO or BSE as output to NAME as selected by CPC PARSER 20264 will place O's onto NAME Concurrently, PARSER 20264 will generate OPREGE to OPCODEREG 20268 to enable transfer of NAME (0- 7) into OPCODEREG 20268. OPCODEREG 20268 is not loaded when PARSER 20264 is parsing Name syllables. The microinstruction input from FUSITT 11012 which controls PARSER 20264 operation also determines whether PARSER 20264 is parsing an SOP or a Name syllable and controls generation of OPREGE.

i- I

J

-310- Operation of NC 10226, which receives Name syllables as address inputs from NAME Bus 20224, has been discussed previously with reference to MEMINT 20212. Name Trap (NT) 20254 is connected from NAME Bus 20224 to receive and capture Name syllables parsed onto NAME Bus 20224 by PARSER 20264. Operation of NT 20254 has been also previously discussed with reference to MEMINT.

b.b.b. Fetch- Unit-DBsiatch Table- -l-i--k.ecgte e Unit-Dispach-TabIe t -20-266-and--Ope-ation r Code--Reqstt;- -2B-2- As previously described, CS 10110 is a multiple language machine. Each program written in a high level c user language is compiled into a corresponding S-Language l program containing S-Language Instructions referred to as SOPs. CS 10110 provides a set or dialect, of microcode instructions, referred to as S-Interpreters (SINTs) for each S-Language. SINTs interpret SOPs to provide ,corresponding sequences of microinstructions for detailed control of CS 10110 operations. CS 10110's SINTs for FU 10120 and EU 10122 operations are stored, respectively, .1 in FUSITT 11012 and in a corresponding control store memory in EU 10122, described in a following description of EU 10122. Each SINT comprises one or more sequences of microinstructions, each sequence of microinstructions corresponding to 4 particular SOP in a particular S- Language dialecto Fetch Unit S-Interpreter Dispatch 4c.

-311- Table (FUSDT) 11010 and Execute Unit S-Interpreter Dispatch Table (EUSDT) 20266 contain an S-Interpreter Dispatcher (SD) fgr each S-Language dialect. Each SD is comprised of a set of SD Pointers (SDPs) wherein each SDP in a particular SD corresponds to a particular SOP of that SD dialect. Each SDP is an address pointing to a location, in FUSITT 11012 or EUSITT, of the start of the corresponding sequence of microinstructions for o interpreting the SOP corresponding to that SDP. As will be described further below, SOPs received and stored in OPCODEREG 20268 are used to generate addresses into FUSDT I t 11010 and EUSDT 20266 to select corresponding SDPs.

St, 4 Those SDPs are then provided to FUSITT 11012 through ADR 20202, or to EUSITT through EUDIS Bus 20206, to select corresponding sequences of microinstructions from FUSITT 'j 11012 and EUSITT.

Referring to Fig. 242, a more detailed block diagram of OPCODEREG 20268, FUSDT 11010, and EUSDT 20266 is shown. As shown therein, OPCODEREG 20268 is comprised of Op-Code Latch (LOPCODE) 24210, Dialect Register (RDIAL) 24212, Load Address Register (LADDR) 24214, and Fetch Unit Dispatch Encoder (FUDISENC) 24216. Data inputs of LOPCODE 24010 are connected from NAME Bus 20224 to receive SOPs parsed from INSTB 20262. Load inputs of RDIAL 24212 are connected from Bits 28 to 31 of JPD Bus 10142. Outputs of LOPCODE 24210, RDIAL 24212 and LADDR 24214 are connected to inputs of FUDISENC 24216. Outputs of FUDISENC 24216 are connected to address inputs of FUSDT 11010 and EUSDT 20266.

-312- FUSDT 11010 is comprised of Fetch Unit Dispatch File (FUDISF) 24218 and Algorithm File (AF) 24220. Address inputs of FUDISF 24218 and AF 24220 are connected, as previously described, from address outputs of FUDISENC 24216. Data load inputs of FUDISF 24218 and AF 24220 are connected from, respectively, Bits 10 to 15 and Bits 16 to 31 of JPD Bus 10142. SDP outputs of FUDISF 24218 and AF 24220 are connected to ADR Buses 20202.

EUSDT 20266 is comprised of Execute Unit Dispatch e 10 File (EUDISF) 24222 and Execute Unit Dispatch Selector (EUDISS) 24224. Address inputs of EUDISF 24222 a~e, as :described above, connected from outputs of FUDISENC 24216. Data load inputs of EUDISF 24222 are connected from Bits 20 to 31 of JPD Bus 10142. Inputs of EUDISS 24224 are connected from SDP output of EUDISF 24222, from t r Bits 20 to 31 of JPD Bus 10142, and from Microcode Literal (mLIT) output of FUSITT 11012. SDP outputs of EUDISS 24224 are connected to EUDIS Bus 20206.

As previously described, OPCODEREG 20268 provides addresses, generated from SOPs loaded into OPCODEREG 20268, to FUSDT 11010 and EUSDT 20266 to select SDPs to be provided as address inputs to FUSITT 11012 and SA EUSITT. LOPCODE 24210 receives and stores eight bit SOPs parsed from INSTB 20262 as described above. OPCODEREG 20268 also provides addresses to FUSDT 11010 and EUSDT 20266 to load FUSDT 11010 and EUSDT 20266 with SDs for S- Language dialects currently being utilized by CS 10110.

LOPCODE 24210 and RDIAL 24212, as described below, provide addresses to FUSDT 11010 and EUSDT 20266 when translating SOPs to SDPs and ADDR 24214 provides I 1 -313-' addresses when FUSDT 11010 and EUSDT 20266 are being loaded with SDs.

Referring first to LADDR 24214, LADDR 24214 has an eight bit counter. Addresses are provided to FUSDT 11010 and EUSDT 20266 from LADDR 24214 only when FUSDT 11010 and EUSDT 20266 are being loaded with SDs, that is groups of SDPs for S-Language dialects currently being utilized by CS 10110. During this operation, output of LADDR S. 24214 is enabled to FUSDT 11010 and EUSDT 20266 by 1io microcode control signals (not shown for clarity of presentation) from FUSITT 11012. Selection between #"0:40 FUDISF 24218, AF 24220, and EUDISF 24222 to receive S* addresses is similarly provided by microinstruction enable signals (also not shown for clarity of presentation) provided from FUSITT 11012. These FUSDT t 11010 and EUSDT 20266 address enable inputs may select, t at any time, any or all of FUDISF 24218, AF 24220, or EUDSF 24222 to receive address inputs. SDPs to be loaded into FUDISF 24218, AF 24220, and EUDISF 24222 are provided, respectively, from Bits 10 to 15 (JPD(10-15)), Bits 16 to 31 (JPD(16-31)), and Bits 20 to 31 31)) of JPD Bus 10142. Address contents of LADDR 24214 are successively incremented by one as successive SDPs are loaded into FUSDT 11010 and EUSDT 20266.

Incrementing of LADDR 24214 is, again, controlled by microinstruction control inputs from FUSITT 11012.

Address inputs to FUSDT 11010 and EUSDT 20266 durng interpretation of SOPs are provided from LOPCODE 24210 I -314and RDIAL 24212. LOPCODE 24210 is a register counter having, as described above, data inputs connected from NAME Bus 20224 to receive SOPs from PARSER 20264. In a first mode, LOPCODE 24210 may operate as a latch, loaded with one SOP at a time from output of PARSER 20264. In a second mode, LOPCODE 24210 operates as a clock register to receive successive eight bit inputs from low order eight bits of NAME Bus 20224 (NAME(8-15)). Loading of LOPCODE 24210 is contolled by microinstruction control outputs (not shown for clarity of presentation) from *FUSITT 11012.

As will be described further below, eight bit SOPs t 4 stored in LOPCODE 24210 are concatenated with the output of RDIAL 24212 to provide addresses to FUSDT 11010 and EUSDT 20266 to select SDPs corresponding to particular C SOPs. That portion of these addresses provided from LOPCODE 24210, that is the eight bit SOPs, selects particular SDPs within a particular 'SD. Particular SDs are selected by that portion of these addresses which is provided from the contents of RDIAL 24212.

RDIAL 24212 receives and stores four bit Dialect 64.' S Codes indicating the particular S-Language dialect currently being used by CS 10110 and executing the SOPs of a user's program. These four bit Dialect Codes are provided from JPD Bus 10142, as JPD (28-31). Loading of RDIAL 24212 with four bit Dialect Codes is controlled by microinstruction control signals provided form FUSITT i 11012 (not shown for clarity of presentation).

Four bit Dialect Codes in RDIAL 24212 define partitions in FUDISF 24218, AF 24220 and EUDISF 24222.

4*

W-'

-315- Each partition contains SDPs for a different S-Language dialect, that is contains a different SD. FUDISF 24218, AF 24220 and EUDISF 24222 may contain, for example, eight 128 word partitions or four 256 word partitions. A single bit of Dialect Code, for example Bit 3, defines whether FUDISF 24218, AF 24220, and EUDISF 24222 contain our or eight partitions. If FUSDT 11010 and EUSDT 20266 S ,contain four partitions, the two most significant bits of address into FUSDT 11010 and EUSDT 20266 are provided r' 10 from Dilect Code Bits 1 and 2 and determine which *e partition is addressed. The lower order eight bits of +tre.: address are provided from LOPCODE 24210 and determine r, which word in a selected partition is addressed. If FUSDT 11010 and EUSDT 20266 contain eight partitions, the three most significant bits of address into FUSDT 11010 and EUSDT 20266 ar provided from Bits 0 to 2 of Dialect Code, to select a particular partition, and the lower seven bits of address are provided from LOPCODE 24210 to t select a particular word in the selected partition.

As described above, LOPCODE 24210 eight oit output and RDIAL 24212's four bit output are concatenated together, through FUDISENC 24216, to provide a ten bit address input to FUSDT 11010 and EUSDT 20266. FUDISENC 24216 is an encoding circuit and will be described further below with reference to FUDISF 24218. As previously described, selection of FUDISF 24218, AF 24220, and EUDISF 24222 to receive address inputs from RDIAL 24212 and LOPCODE 24210 is controlled by S- \I i -316microinstruction control enable inputs provided from FUSITT 11012 (not shown for clarity of presentation).

Referring to FUSDT 11010, both FUDISF 24218 and AF 24220 provide SDPs to FUSITT 11012, but do so for differing purposes. In general, microinstruction control operations may be regarded as falling into two classes.

First, there are those microinstruction operations which .are generic, that is general in nature and used by or applying to a broad variety of SOPs of a particular 10 dialect or even of many dialects. An example of this .0 class of microinstuction operation is fetches of operand values. FUDISF 24218 provides SDPs for this class of microinstruction operations. As described below, FUDISF 24218 is a fast access memory allowing a single microinstruction control output of FUSITT 11012 to parse t an SOP from INSTB 20262 into LOPCODE 24210, and a Stt r, corresponding SDP to be provided from FUDISF 24218. That is, an SOP of this generic class may be parsed from INSTB St "20262 and a corresponding SDP provided from FUDISF 24218 during a single system clock cycle. Operation of FUDISF 24218 thereby enhances speed of operation of JP 10114, in particular at the beginning of execution of new SOPs.

The second class of microinstruction operations are those specific to particular SINTs or to particular groups of SINTs. These groups of SINTs may reside entirely within a particular dialect, for example FORTRAN, or may exist within one or more dialects. SDPs for this class of microinstruction operation are provided by AF 24220. As described further below, AF 24220 is slower than FUDISF 24218, but is larger. In general, AF L ~tl ~r- -317a 44 4 4 a C ai.

@4 Sa ad a.

a 44 24220 contains SDPs of microinstruction sequences specific to particular SINTs. In general, generic microinstruction operations are performed before those operations specific to particular SINTs, so that SDPs are required from AF 24220 at a later time than those from FUDISF 24218. SDPs for specific SINT operations may therefore be provided from lower speed AF 24220 without a penalty in speed of execution of SOPs.

Referring again to FUDISF 24218, FUDISF 24218 is a 10 1,024 word by 6 bit fast access by polar memory. Each word contained therein, as described above, is an SDP, or address to start of a corresponding sequence of microinstructions in FUSITT 11012. As will be described further below, FUSITT is an 8K (8192) word memory. SDPs 15 provided by FUDISF 24218 are each, as described above, 6 bits wide and may thus address a limited, 32 word area of FUSITT 11012's address space. FUDISF 24218 is enabled to provide SDPs to FUSITT 11012 by microinstruction control signals (not shown for clarity of presentation) from FUSITT 11012. FUDISF 24218 six bit SDPs are encoded by FUDISENC 24219 to address FUSITT 11012 address space in increments of 4 microinstructions, that is in increments of 4 address locations. FUDISF 24218 SDPs thereby address 4 microinstructions at a time from FUSITT 11012's microinstruction sequences. As will be described'further below, mPC 20276 generates successive microinstruction addresses to FUSITT 11012 to select successive microinstructions of a sequence following an initial #3a4 a i a a, a 4 4 i: Tj 4

L

Iri ii -318microinstruction selected by an SDP from FUSDT 11010. An FUDISF 24218 SDP will thereby select the first microinstruction of a 4 microinstruction block, and mPC 20276 will select the following 3 microinstructions of that 4 microinstruction sequence. A 4 microinstruction sequence may therefore be executed in line, or sequentially, for each FUDISF 24218 SDP provided in response to a generic SOP. FUDISENC 24219 encodes FUDISF i 24218 six bit SDPs to select these 4 microinstruction S,0 10 sequences so that the least significant bit of these SDPs occupies the 24 bit of FUSITT 11012 address inputs, and 'so on. The two least significant bits of an FUSITT 11012 Saddress, or SDP, provided from FUDISF 24218 are forced to 0 while the ninth and higher bits may be hard-wired to define any particular block of 128 addresses in FUSITT 11012. This hard-wiring of the most significant bits of S FUSITT 11012 addresses from FUDISF 24218 allows a set of generic microinstruction sequences selected by FUDISF 24218 to be located as desired within FUSITT 11012's address space. FUDISENC 24219 is comprised of a set of driver gates.

As previously described, SDPs for generic microinstructions currently being utilized by CS 10110 in executing user's programs are written into FUDISF 24218 from Bits 10 to 15 of JPD Bus 10142 (JPD(10-15)).

Addresses for loading SDPs into FUDISF 24218 are provided, as previously described, from LADDR 24214.

LADDR 24214 is enabled to provide load addresses, and FUDISF 24218 is enabled to be written into, by -319microinstruction control signals (not shown for clarity of presentation) provided from FUSITT 11012.

Referring to AF 24220, as previously described AF 24220 is of larger capacity than FUDISF 24218, but has slower access time. AF 24220 is a 1,024 word by 15 bit memory. In general, 2 clock cycles are required to obtain a DSP from AF 24220. During first clock cycle, an *c, 1 SOP is loaded into LOPCODE 24210 and, during second clock S. cycle, AF 24220 is addressed to provide a corresponding SDP. SDPs provided by AF 24220 are each 15 bits in width cc 4* and thus capable of addressing a larger address space t^ than that of FUSITT 11012. As previously described, t FUSITT 11012 is' an 8K word memory. If FUSITT 11012 is addressed by an AF 24220 SDP referring to an address location outside of FUSITT 11012's address space, FUSITT 11012 will generate a microinstruction Not In Control S Store output to EVENT 20284 as described further below.

t An AF 24220 SDP resulting in this event will then be used 'to address certain microinstruction sequences stored in MEM 10112. These microinstructions will then be executed from MEM 10112, rather than from FUSDT 11010. This operation allows certain microinstruction sequences, for 444949 example rarely used microinstruction sequences, to remain in MEM 10112, thus freeing AF 24220 and FUSITT 11012's address spaces from more frequently used SOPs.

As previously described AF 24220 is loaded, with SDPs, for SINTs currently being used by CS 10110 in executing user's programs, from Bits 16-31 of JPD Bus 10142 (JPD(16--31)). Also as previously discussed, addresses to load SDPs into AF 24220 are provided from i t -320t...I Atc j" i i£ t tr i'4

A.

IL at I LADDR 24214. LADDR 24214 is enabled to provide load addresses and AF 24220 to receive SDPs, by microinstruction control signals (not shown for clarity of presentation) provided from FUSITT 11012.

Referring finally to EUSDT 20266, SDPs may be provided to EU 10122 from 3 sources. EU 10122 SDPs may be provided from EUDISF 24222, from JPD Bus 10142 or from literal fields of microinstructions provided from FUSITT 11012. EUDISF 24222's SDPs are each 12 bits in width and comprise 9 bits of address into EUSITT and 3 bits of operand format information.

EUDISF 24222-is 1,024 word by 12 bit memory. As previously described addresses to read SDPs from EUDISF 24222 are provided from OPCODEREG 20268 by concatenating 15 a 4 bit Dialect Code from RDIAL 24212 and an 8 bit SOP from LOPCODE 24210. SDPs provided by EUDISF 24222 are provided as a first input to EUDISS 24224.

EUDISS 24224 is a multiplexer. As just described, a first input of EUDISS 24224 are SDPs from EUDISF 24222.

A second 12 bit input of EUDISS 24224 is provided from Bits 20 to 31 of JPD Bus 10142 (JPD(20-31)). A third input of EUDISS 24224 is a 12 bit input provided from a literal field of an FUSITT 11012 microinstruction output. EUDISS 20224 selects one of these 3 inputs to be transferred on EUDIS Bus 20206 to be provided as an execute unit SDP to EUSITT. Selection between EUDISS 20224's inputs is provided by microinstruction control signals (not shown for clarity of presentation) provided from FUSITT 11012.

-321- As previously described, EUDISF 24222 is loaded, with SDPs for S-Language dialects currently being used by CS 10110, from Bits 20 to 31 of JPD Bus 10142 (JPD(20-31)).

Addresses to load SDPs into EUDISF 24222 are provided, as previously described, from LADDR 20214. FUSITT 11012 provides enable signals (not shown for clarity of presentation) to LADDR 24214 and EUDISF 24222 to enable writing of SDPs into EUDISF 24222.

The structure and operation of FUCTL 20214 circuitry 4 a 10 for fetching and parsing SINs from MEM 10112 to provide SIt Name syllables and SOPs, and for interpreting SOP to P provide SDPs to FUSITT 11012 and EUSITT from FUSDT 11010 and EUSDT 20266, have been described above. As described above, SDPs provided by FUSDT 11010 and EUSDT 20266 are initial, or starting, addresses pointing to first microinstructions of sequences of microinstructions.

Addresses for microinstructions following those initial Sf microinstructions are provided by FUCTL 20214's next i address generator circuitry which may include mPC 20276, BRCASE 20278, REPCTR 20280 and PNREG 20282, EVENT 20284 and SITTNAG 20286. mPC 20276, BRCASE 20278, REPCTR 20280 and PNREG 20282, and SITTNAG 20286 are primarily C l concerned with generation of next addresses during Sexecution of microinstruction sequences in response to SOPs and will be described next below. EVENT 20284 and other portions of FUCTL 20214's circuitry are more concerned with generation of microinstruction sequences with regard to CS 10110's internal mechanisms operations 1-I- -322and will be described in a later description. EU 10122 also includes next address generation circuitry and this circuitry will be described in a following description of EU 10122.

c.c.c. Next-Address-Generator-24-a- 0 -Fgiag...243- As stated above, in FU 10120 first, or initial, microinstuctions of microinstruction sequences for interpreting SOPs are provided by FUSDT 11010-.

S* 10 Subsequent addresses of microinstructions within these "f'i sequences are, in general, provided by mPC 20276 and BRCASE 20278. mPC 20276, as described further below, provides sequential addresses for selecting sequential microinstructions of microinstruction sequences. BRCASE 20278 provides addresses for selecting microinstructions when a microinstruction Branch or microinstruction Case operation is required. REPCTR 20280 and PNREG 20282 provide addresses for writing, or loading, of Smicroinstruction sequences into FUSITT 11012. Other portions of FUCTL 20214 circuitry, for example EVENT 20284, provides microinstruction sequence selection addresses to select microinstruction sequences for controlling operation of CS 10110's internal mechanisms.

SITTNAS 20286 selects between these microinstruction address sources to provide to FUSITT 11012 those addresses required to select micrQonstructions of the operation to be currently executed by CS 10110.

L^ -323- Referring to Fig. 243, a partial block diagram of FU 10120's Next Address Generator (NAG) 24310 is shown. In addition to FUSDT 11010, NAG 24310 includes mPC 20276, BRCASE 20278, EVENT 20284, REPCTR 20280 and PNREG 20282, a part of RCWS 10358, and SITTNAS 20286. EVENT 20284 is, as described above, primarily concerned with execution of microinstruction sequences for controlling CS 10110 internal mechanisms. EVENT 20284 as shown herein only to e, illustrate its relationships to other portions of NAG 24310. EVENT 20284 will be described further in a following description of FUCTL 20214's circuitry' controlling CS 10110's internal mechanisms. Similarly, S/ operation of RCWS 10358 will be described in part in the present description of NAG 24310, and in part in a following description of control of CS 10110's internal Smechanisms.

Se. Referring first to NAG 24310's structure, interconnections of FUSDT 11010, RCWS 10358, mPC 20276, BRCASE 20278, REPCTR 20280, PNREG 20282, EVENT 20284, and SITTNAS 20286 have been previously described with reference to Fig.. 202. NAG 24310's structure will be described below only wherein Fig. 243 differs from Pig.

202.

Referring first to SITTNAS 20286, SITTNAS 20286 is shown as comprised of EVENT Gate (EVNTGT) 24310 and Next Address Select Multiplexer (NASMUX) 24312. NASMUX 24312 is comprised of NAS Multiplexer A (NASMUXA) 24314, NASMUXB 24316, NASMUXC 24318, and NASMUXD 24320. Outputs -324of EVNTGT 24310 and NASMUXA 2431.4 to NASMUXD 24320 are ORed to CSADR 20204 to provide microinstruction selection addresses to FUSITT 11012.

ADR 20202 is shown in Fig. 243 as comprised of nine buses, Address A (ADRA) Bus 24322 to Address I (ADRI) Bus 24338. Output of EVENT 20284 is connected to input of EVNTGT 24310 by ADRA Bus 24322. Outputs of REPCTR 20280 and PNREG 20282 and output of AF 24220 are connected to inputs of NASMUXA 24314 by, respectively, ADRB Bus 24324 and ADRC Bus 24326. Outputs of RCWS 10358 and FUDISENC 24219 are connected to inputs of NASMUXB 24316 by, respectively, ADRD Bus 24328 and ADRE Bus 24330. Outputs of BRCASE 20278 and second output of mPC 20276 are connected to inputs of NASMUXC 24318 by, respectively, ADRF Bus 24332 and ADRG Bus 24334. Second output of mPC 20276 and JAM output of NC 10226 are connected to inputs of NASMUXD 24320 by, respectively, ADRH Bus 24336 and ADRI Bus 24338. ADR 20202 thus comprises a set of buses connecting microinstruction address sources to inputs of SITTNAS 20286.

Referring to mPC 20276, mPC 20276 is comprised of SMicro-Program Counter Counter (mPCC) 24340 and Micro- Program Counter Arithmetic and Logic Unit (mPCALU) 24342. Data input of mPCC 24340 is connected from CSADR Bus 20204. Output of mPCC 24340 is connected to a first input of mPCALU 24342 and is mPC 20276's third ovtput to BRCASE 20278. Second input of mPCALU 24342 is a fifteen binary number set, for example by hard-wiring, t o be first output, to RCWS 10358, and mPC 20276's second output, to inputs of NASMUXC 24318 and NASMUXD 24320.

BRCASE 20278 is shown in Fig. 243 as comprising Mask and Shift Multiplexer (MSMUX) 24344, Case Mask and Shift Logic (CASEMS) 24346, Branch and Case Multiplexer (BCMUX) 24348 and Branch and Case Arithmetic and Logic Unit (BCALU) 24350. A first input of MSMUX 24344 (AONBC, not previously shown) is connected from output of AONGRF I 10 20232. A second input of MSMUX 24344 (OFFMUXR, not S previously shown) is connected from output of OFFMUXR S 23812. Output of MSMUX 24344 is connected to input S CASEMS 24346, and output of CASEMS 24346 is connected to a first input of BCMUX 24348. A second input of BCMUX 24348, BLIT is connected from a literal field output of S, FUSITT 11012's microinstruction output. Output of BCMUX 24348 and third output of mPC 20276, from output of mPCC 24340, are connected, respectively, to first and second 4 inputs of BCALU 24350. Output of BCALU 24350 comprises BRCASE 20278 outputs to NASMUXC 24318.

An address to select a next microinstruction may be provided to FUSITT 11012 by SITTNAS 20286 from any of ,eight sources. First source is output of mPC 20276.

P. Output of mPC 20276 is referred to as Micro-Program Count Plus 1 (mPC+1) and is fifteen bits of address. Second source is from EVENT 20284 and is comprised of five bits of address. Third source is output of FUDISP 24218 and FUDISENC 24219 and, as previously described, is comprised L 4 FI l. i ib 7- -326r 44 .r 41 I Ot 4 *4 CI 44 of six bits of address. Fourth source is output of AF 24220 and, as previously described, is comprised of fifteen bits of address. Fifth source is output of BRCASE 20278. Output of BRCASE 20278 is referred to as Branch and Case Address (BRCASEADR) and comprises fifteen bits of address. Sixth source is an output of RCWS 10358. Output of RCWS 10358 is referred to as RCWS Address (RCWSADR) and is comprised of fifteen bits of address. Seventh source is REPCTR 20280 and PNREG 20282 10 whose outputs (REPPN) together comprise fifteen bits of address. Finally, eighth source is JAM input from NC 10226, which comprises five bits of address. These address sources differ in number of bits of address that they provide, but a microinstruction address gated onto CSADR Bus 20202 by SITTNAS 20286 always comprises fifteen bits of address. If a particular source applies fewer than fifteen bits, that address is extended to fifteen bits by SITTNAS 20286. In general, extension of address bits may be performed by hard-wiring of additional address input bits to SITTNAS 20286 from each of these sources and will be described further below.

Referring to mPC 20276, mPCC 24340 is a fifteen bit register and mPCALU 24342 is a fifteen bit ALU. mPCC 24340 is, as described above, connected from CSADR Bus 20204 and is sequentially loaded with a microinstruction address currently being presented to FUSITT 11012. mPCC 24340 will thus contain the address of the currently executing microinstruction. mPCALU 24342 is dedicated to j -327- 4* a *n a, a ii ai I a, 4*4a incrementing the address contained in mPCC 24340 by one.

mPC+1 output of mPCALU 24342 will thereby always be address of next sequential microinstruction. mPC+1 is, as described above, a fifteen bit address and is thus not extended in SITTNAS 20286.

Referring to BRCASE 20278, as described above BRCASE 20278 provides next microinstruction addresses for mPC 20276 Relative Branches and for Case Branches. Next microinstruction addresses for microprogram Relative Branches and for Case Branches are both generated as addresses relative to address of currently executing microinstruction as stored in mPCC 24340, but differ in the manner in which these relative addresses are generated. Considering first Case Branches, Case Branch 15 addresses relative to a currently executing microinstruction address are generated, in part, by MSMUX 24344 and CASEMS 24346. As described above, MSMUX 24344 which is a multiplexer receives two inputs. First input is AONBC from output of AONGRF 20232 and second input is OFFMUXR from output of OFFMUXR 23812. Each of these inputs is eight bits, or one byte, in width. Acting under control of microinstruction output from FUSITT 11012, MSMUX 24344 selects either input AONBC or input OFFMUXR as an eight bit output to input of CASEMS 24346.

CASEMS 24346 is a Mask and Shift circuit, similar in structure and operation to that of FIU 20116 but operating upon bytes rather than thirty-two bit words.

CASEMS 24346, operating under microinstruction control from FUSITT 11012, manipulates eight bit input from MSMUX 24344 by masking and shifting to provide eight bit Case *q as1 *4 5 4r~ I a -328- Value (CASEVAL) output to BCMUX 24348. CASEVAL represents a microinstruction address displacement relative to address of a currently executing microinstruction and, being an eight bit number, may express a displacement of 0 to 255 address locations in FUSITT 11012.

BCMUX 24348 is an eight bit multiplexer, similar in structure and operation to MSMUX 24344, and is controlled by microinstruction inputs provided from FUSITT 11012.

In executing a case operation, BCMUX 24348 selects CASEVAL input to MCMUX 24348's output to first input of BCALU 24350. BCALU 24350 is a sixteen bit arithmetic and logic unit. Second input of BCALU 24350 is fifteen bit address of currently executing microinstruction from mPCC 24340. BCALU 24350 operates under microinstruction I control provided from FUSITT 11012 and, in executing a i Case operation, adds CASEVAL to the address of a currently executing microinstruction. During a Case operation, carry input of B.ALU 24350 is forced, by microinstruction control from FUSITT 11012, to one so that BCALU 24350's second input is effectively mPC+1, or i address of currently executing microinstruction plus 1.

Output BRCASEADR of BCALU 24350 will thereby be fifteen bit Case address which is between one and 256 FUSITT 11012 address locations higher than the address location Sof the currently executing microinstruction. The actual case value address displacement from the address of the currently executing microinstruction is determined by -329either input AONBC or input OFFMUXR to MSMUX 24344, and these mask and shift operations are performed by CASEMS 24346.

Case operations as described above may be used, for example, in interpreting and manipulating CS 10110 table entries. For example, Name Table Entries of Name Tables 10350 contain flag fields carrying information regarding certain operations to be peformed in resolving and evaluating those Name Table Entries. These operations 10 may be implemented as Case Branches in microinstruction sequences for resolving and evaluating those Name Table t Entries. In the present example, during resolve of a S•.Name Table Entry the microinstruction sequence for performing that resolve may direct a byte of that Name Table Entry's flag field to be read from AONGRF 20232, or :tt" OFFMUXR 23812, and through MSMUX 24344 to CASEMS 24346.

That microinstruction sequence will then direct CASEMS 24346 to shift and mask that flag field byte to provide a I CASEVAL. That CASEVAL will have a value dependent upon the flags within that flag field byte and, when added to mPC+1, will provide a FUSITT 11012 microinstruction address for a microinstruction sequence for handling that Name Table Entry in accordance with those flag bits.

As described above, BRCASE 20278 may also generate microinstruction addresses for Branches occurring within execution of a given microinstruction sequence. In this case, microinstruction control signals from FUSITT 11012 i direct BCMUX 24348 to select BCMUX 24348's second input y:*s -330as output to BCALU 24350. BCMUX 24348's second input is Branch Literal (BLIT). As described above, BLIT is provided from a literal field of a microinstruction word from FOSITT 11012's microinstruction output. BLIT output of BCMUX 24348 is added to address of currently executing microinstruction from mPCC 24340, and BCALU 24350, to provide fifteen bit BRCASEADR of a microinstruction address branched to from the address of.the currently executing microinstruction. BRCASEADR may represent, for example, any of four Branch Operations. Possible Branch Operations are: first, a Conditional Short Branch; second, a Conditional Short Call; third, a Long Go To; and, fourth, a Long Call. In each of these possible Branch Operations, BLIT is treated as the twos complement of the desired branch value, that is the microinstruction S: address offset relative to the address of the currently executing microinstruction. BLIT field may therefore be, effectively, added to or subtracted from the address of t, the currently executing microinstruction, to provide a microinstruction address having a positive or negative displacement from the address of the currently executing I microinstruction. In a Conditional Short Brancsh or a Conditional Short Call, the fourteen bit literal field is a sign extended eight bit number. Both Conditional Short Branch and Conditional Short Call microinstruction addresses may therefore point to an address within a range of +127 to -128 FUSITT 11012 address locations of the address of the currently executing microinstruction.

In the case of a Long Go To or Long Call, the BLIT field L_ 1 l--i/i l -331is a fourteen bit number representing displacement relative to the address of the currently executing microinstruction. BRCASEADR may, in these cases, represent a FUSITT 11012 microinstruction address within a range of +8191 to -8192 FUSITT 11012 address locations of the address of the currently executing microinstruction. BRCASE 20278 thereby provides FU 10120 'with capability of executing a full range of microinstruction sequence Case and Branch operations.

Referring to RCWS 10358, as previously described RCWS 10358 stores information regarding microinstruction sequences whose execution has been halted. RCWS 10358 f allows execution of those microinstruction sequences to be resumed at a later time. A return control word (RCW) may be written onto RCWS 10358 during any microinstruction sequence that issues a Call to another Smicroinstruction sequence. The calling microinstruction S sequence may, for example, be aborted to service an event, as described further in a following description, or may result in a Jam. A Jam is a call for a microinstruction sequence which is forced by operation of 4i CS 10110.hardware, rather than by a microinstruction sequence. RCWS 10358 operation with regard to CS 10110's internal mechanisms will be described in a following description of EVENT 20284, STATE 20294, and MCW1 20290 and MCWO 20292. For puposes of the present discussioni, that portion of a RCW concerned with interpretation of SOPs contains, first, certan state information from FUSITT 11012 and, second, a return address into FUSITT -332- 11012. State that FUSITT 11012 state is provided from STATE 20294, as described below, and that portion of a RCW containing FUSITT 11012 state information will be described in a following description. Microinstruction address portions of RCWs are provided from output of mPCALU 24342. This microinstruction address is the address of the microinstruction to which FU 10120 is to return upon return from a Call, Event, or Jam. Upon *o occurrence of a Call or Jam, the microinstruction return 10 address is mPC+1, that is the address of the microinstruction after the microinstruction issuing the SCall or Return. For aborted microinstruction sequences, S.4 the microinstruction return address is mPC, that is the address of the microinstruction executing at the time abort occurs.

Upon return from a call, service of an event, or service of a jam, FU 10120 state flag portion of RCW is loaded into STATE 20294. Microinstruction return address Ott is provided by RCWS 10358 as fifteen bit RCWSADR to SITTNAS 20286 and is gated onto CSADR 20204. RCWSADR is provided to FUSITT 11012 to select the next microinstruction and is loaded into mPCC 24340 from CSADR 20204.

SAs previously described, RCWS 10358 is connected to JPD Bus 10142 by a bi-directional bus. RCWs may be written into RCWS 10358 from JPD Bus 10142, or read from RCWS 10358 to JPD Bus 10142. The fifteen bit next microinstruction address portion, and the single bit FUSITT 11012 state portion of RCW is written from or read -333to Bits 16 to 31 of JPD Bus 10142. FU 10120 may write Present Bottom RCW or Previous RCW into RCWS 10358 from JPD Bus 10142 and may read Present Bottom RCW, or Previous RCW, or another selected RCW, onto JPD Bus 10142. RCWS 10358 thereby provides a means for storing and returning microinstruction addresses of microinstruction sequences whose execution has been suspended, and a means for writing and reading microinstruction address, and FUSITT 11012 state flags, 10 from and to JPD Bus 10142.

As previously described, REPCTR 20280 and PNREG 20282 provide microinstruction addresses for writing of .0 microinstructions into FUSITT 11012. REPCTR 20280 is an eight bit counter and PNREG 20282 is a seven bit register. Eight bit output of REPCTR 20280 is left concatenated with seven bit output of PNREG 20282 to provide fifteen bit microinstruction addresses REPPN.

That is, REPCTR 20280 provides the eight low order bits 4 of microinstruction address while PNREG 20282 provides the seven most significant bits of address.

REPCTR may be loaded from Bits 24-31 of JPD BUS 10142, and may be read to Bits 24-31 of JPD Bus 10142.

In addition, the eight bits of microinstruction address in REPCTR 20280 may be incremented or decremented as microinstructions are written into FUSITT 11012.

As described above, PNREG 20282 contains the seven most significant bits of microinstruction address. These address bits may be written into PNREG 20282 from Bits 17-23 of JPD Bus 10142. Contents of PNREG 20282 may not, in LP 1 1 -334general, be read to JPD Bus 10142 and may not be incremented or decremented.

Referring to JAM input to SITTNAS 20286 from NC 10226, certain Name evaluate or resolve operations may result in jams. A Jam functions as a call to microinstruction sequences for servicing Jams and are forced by FU 10120 hardware circuitry involved in Name syllable evaluates and resolves.

JAM input to SITTNAS 20286 is comprised of six Jam 10 address bits. Three bits are provided by NC 10226 and three bits are provided from FUSITT 11012's microinstruction output as part of microinstruction sequences for correcting Name syllable evaluates and resolves. The three bits of address from NC 10226 form the most significant three bits of JAM address. One of these bits gates JAM address onto CSADR Bus 20204 and is 4 thus not a true address bit. Output of FUSITT 11012 provides the three least significant bits of JAM address ,and specifies the particular microinstruction sequence required to service the particular Jam which has occurred. Therefore, during Name evaluate or resolves, the microinstruction sequences provided by FUSITT 11012 4 t to perform Name evaluates or resolves specifies what microinstruction sequences are to be initiated if a Jam occurs. The three bits of JAM address provided by NC 10226 determine, first, that a Jam has occurred and, second, provide two bits of address which, in combination with the three bits of address from FUSITT 11012, specify the particular microinstruction sequence for handling 1 1 1 7 -335- 4* 44 4 4i 1 i Il that Jam. JAM address inputs from NC 10226 and from FUSITT 11012 thereby provide six of the fifteen bits of JAM address. The remaining nine bits of JAM address are provided, for example, by hard-wired inputs to NASMUXD 24320. These hard-wired address bits force JAM address to address FUSITT 11012 in blocks of 4 microinstruction addresses, in a manner similar to address inputs to FUDISF 24218 and FUDISENC 24219.

Address inputs provided to SITTNAS 20286 from FUSDT 11010 have been previously described with respect to description of FUCTL 20214 fetch, parse, and dispatch operations. Address inputs provided by EVENT 20284 will be described in a following description of FUCTL 20214's operations with regard to CS 10110's internal mechanisms.

Referring finally to SITTNAS 20286, as previously described SITTNAS 20286 is comprised of EVNTGT 24310 and NASMUX 24312. Inputs are provided to NASMUX 24312, as described above, from FUSDT 11010, mPC 20276, BRCASE 20278, RCWS 10358, REPCTR 20280 and PNREG 20282, and by JAM input. These inputs are, in general, provided with regard to FUCTL 20214's operations in fetching, parsing, and interpreting SOPs and Name syllables. These operations are thereby primarily directly concerned with execution of user's programs, that is the execution of sequences of SINs. NASMUX 24312 selects between these inputs and transfers selected address inputs onto CSADR 20204 as microinstruction addresses to FUSITT 11012 under microinstruction control from microinstruction outputs of FUSITT 11012. Microinstruction address outputs are -336provided to SITTNAS 20286 from EVENT 20284 in response to Events, described further below, occuring in CS 10110's operations in executing user's programs. These microinstruction addresses from EVENT 20284 are gated onto CSADR 20204, to select appropriate microinstruction sequences, by EVNTGT 24310. EVNTGT 24310 is separated from NASMUX 24312 to allow EVNTGT 24310 to over-ride NASMUX 24312 and provide microinstruction address to EVENT 20284 while NASMUX 24312 is inhibited due'to 10 occurrence of certain Events. These Events are, in Sgeneral, associated with operation of CS 10110's internal mechanisms and structure and oepration of EVENT 20284, together with STATE 20294, MCW1 20290, and MCWO 20292, and other portions of RCWS 10358, will be described next below.

c.c. UCU T--2O-209i- '*tr Certain porzions of FUCTL 2'0214's Control Circuitry are more directly concernaed with operation of CS 10110's internal mechanizmsm orC example CS 10110 Stack Mechanisms. This circuitry may include STATE 20294, S EVENT 20284, MCW1 20290 and MCWO 20292, portions of RCWS 10358, REG 20288, and Timers 20296. These FUCTL 20214 control elements will be described next below, beginning with STATE 20294.

a.a.a. State-*.ogIc- 20-294 In general, all CS 10110 operations, including execution of microinstructions, are controlled by CS f 5 o n t o CSADR 20204, to select aporaemcontuto -337- 10110's Operating State. CS 10110 has a number of Operating States, hereafter referred to as States, each State being defined by certain operations which may be performed in that State. Each of these States will be described further below. Current State of CS 10110 is indicated by a set of State Flags stored in a set of registers in STATE 20294. Each State is entered from previous State and is exited to a following State. Next State of CS 10110 is detected by random logic' gating 10 distributed throughout CS 10110 to detect certain conditions indicating which State CS 10110 will enter next. Outputs of these Next State Detection* gates are provided as inputs to STATE 20294's registers. A 1 4 particular State register is set and provides a State Flag output when CS 10110 enters the State associated with that particular register. State Flag outputs of STATE 20294's state registers are provided as enable signals throughout CS 10110 to enable initiaton of operations allowed within CS 10110's current State, and to inhibit initiation of operations which are not allowed within CS 10110's current State.

t Certain of CS 10110's States, and associated STATE 20294 State Regis.ters and State Flag outputs, are: MO: the initial State of any microinstrucilon.

State MO is always entered as first data cycle of every microinstruction. During MO, CS 10110's State may not be changed, thus allowing a microinstruction to be arbitrarily aborted and restarted from State MO. In normal execution of microinstructions, State MO is l.

1 -338followed by State M1, described below, that is, State MO is exited to State M1. State MO may be entered from State MO and from State Ml, State AB, State LR, State NR, or State MS, each of which will be described below.

EP: Enable Pause State. State EP is entered when State MO is entered for the first time in a microinstruction. If that microinstruction requests a pause, that microinstruction will force State MO to be Ire-entered for one clock cycle. If State MO lasts more S 0 than one clock cycle, State EP is entered on each extension of State MO unless the extension is a result of a pause request.

S(3) SR: Source GRF State. SR State is active for .one clock cycle wherein SR State register enables loading of a GRF 10354 output register. State SR is re-entered on every State MO cycle except a State MO cycle generated S by a microinstruction requesting extension of State MO.

i 10 When all STATE 20294 State Registers are cleared, DP r, 20218 may set state SR register alone, for purposes of reading from GRF 10354.

Ml: Final state of normal microinstruction execution. State M1 is the exit State of normal S 1 mitSistructaion execution. FUSITT 11012 microinstruction register, described below, is loaded with a next microinstruction upon exit from State MI. In addition, State Ml Flag output of STATE 20294 enables all CS 10110 registers to receive data on their inputs, that is data on inputs of these registers are clocked to outputs of these registers. State M1 may be entered from it State MI, or from State MO, State MW, State MWA, or State

WB.

LA: Load Accumulator Enable State. State LA is entered, upon exit from State M1, by microinstructions which read data from MEM 10112 to OFFMUXR 23812. As previously described, OFFMUXR 23812 serves as a general purpose accumulator for DESP 20210. STATE LA overlaps into execution of next microinstruction, and persists until data is returned from MEM 10112 in response to a request to MEM 10112. When MEM 10112 signals data is t available, by asserting DAVFA, LA State Flag enables loading of data into OFFMUXR 23812. If the next microinstruction references OFFMUXR 238,12, that S< microinstruction execution is deferred until a read to OFFMUXR 23812 is completed, as indicated by CS 10110 exiting from State LA.

RW: Load GRF 10354 Wait State. State RW is t entered from State M1 of microinstructions which read data from MEM 10112 to GRF 10354. RW Fag inhibits S 20 initiation of a next microinstruction, that is prevents entry to State MO, and persists through the CS 10110 Sclock cycle during which data is returned from MEM 10112 in response to a request. State RW initiates Load GRF SEnable State, described below.

LR: Load GRF Enable State. State LR is entered in parallel with State RW, on last clock cycle of RW, and persists for one CS 10110 clock cycle. LR Flag enables writing of MEM 10112 output data into GRF 10354.

pp. -340- 9 S 0 6*a 6 *L 6 6* 6 6 9

S.

0 6 MR: Memory Reference Trailer State. State MR is entered on transition to State MO whenever a previous microinstruction makes a logical or physical address reference to MEM 10112. MR Flag enables recognition of any MEM 10112 reference Events, described below, which may occur. State MR persists for one clock cycle. If an MEM 10112 memory reference Event occurs, that Event forces exit from State MR to States AB and MA, otherwise State MR has no effect upon selection next state.

10 SB: Store Back Enable State. State SB is entered during State MO of a microinstruction following a microinstruction which generated a store back of a result of a EU 10122 operation. SB Flag gates that result to be written into MEM 10112 through JPD Bus 10142.

(10) AB: Microinstruction Abort State. State AB is entered from first MO State after an Event request is recognized, as described in a following description.

State AB may be entered from State MO or from State AB and suppresses an entry into State Ml. If there has been an uncompleted reference to MEM 10112, that is, the reference has not been aborted and data has not returned from MEM 10112, JP 10114 remains in State AB until the MEM 10112 reference is completed. Should an abort have occurred due to a MEM 10112 reference Event, State AB lasts two clock cycles only. As will be described in a following description of EVENT 20284, State MO of a first microinstruction of a Handler for an Event causing an abort is entered from State AB. AB Flag gates the Handler address of the highest priority recognized Event jl itl Fd -341onto CSADR Bus 20204 to select a corresponding Event Handler microinstruction sequence. EVENT 20284 is granted control of CSADR Bus 20204 during all State AB clock cycles.

(11) AR: Microinstruction Abort Reset State. State AR is entered in parallel with first clock cycle of State AB and persists for one clock cyc'e. AR Flag resets various STATE 20294 State Registers when an abort occurs. If there are no uncompleted MEM 10112 references, next State AB clock cycle is the last. On Suncompleted MEM 10112 references, State AR is entered, but State AB remains active until reference is complete.

SShould a higher priority Event request service and be Srecognized while JP 10114 is in State AB, State AR is reentered. State AB will thereby be active for two clock cycles during all honored Event requests.

(12) MA: MEM 10112 Reference Abort. State MA is S entered in parallel with State AB if a MEM 10112 t Ce: reference is aborted, as indicated by asserted ABORT control signal output from MEM 10112. State MA persists for one clock cycle and State AB flag generates a MEM 10112 Reference Abort Flag which, as described below, S' results in a repeat of the MEM 10112 reference. AB Flag Salso resets MEM 10112 Trailer States, described below.

(13) NW: Nano-interrupt Wait State. State NW is entered from State MO of a microinstruction which issues a Nano-interrupt Request to EU 10122 for an EU 10122 operation. FU 10120 remains in State NW until EU 10122 acknowledges that interrupt. Various EU 10122 Events may ~i ,,rfrnei botd sidctd yasre BR -342make requests at this time. State NW is exited into State AB or State Ml.

(14) FM: First Microinstruction of a SIN. State FM is entered in parallel with State MO on first microinstruction of each SIN and persists for one clock cycle. FM Flag inhibits premature use of AF 24220 and enables recognition of SIN Entry Events. State FM is reentered upon return from all aborts taken during State MO of the first microinstruction of an SIN.

'10 (15) SOP: Original Entry to First SIN. State SOP is entered upon entry to State MO of the first microinstruction of an SOP and is exited from upon any exit from that microinstruction. State SOP is entered 0 f only once for each SOP. SOP Flag may be used, for example, for monitoring performance of JP 10114.

(16) EU: EU 10122 Operand Buffer Unavailable. State EU is entered from State MO of a microinstruction which attempts to read data to EU 10122 Operand Buffer, described in a following description, wherein EU 10122 Operand Buffer is full. When a new SOP is entered, three fetches of data from MEM 10112 may be performed before EU 10122 Operand Buffer is full; two fetches will fill EU 10122 Operand Buffer but EU 10122 may take one operand during a second fetch, thereby clearing EU 10122 Operand Buffer space for a third operand.

(17) NR: Long Pipeline Read. Entry into State NR disables overlap of MEM 10112 reads and disables I execution of the next microinstruction. A following microinstruction does not enter State MO until requested Lfdata is returned from MEM 10112. State NR is entered from State NR or from State Ml.

(18) NS: Nonpipeline Store Back. State NS is entered in parallel with State SB whenever a microinstruction requesting a pipeline store back, or a write to MEM 10112, occurs. State NS flag generates entry into State MO of a following microinstruction upon exit from State SB.

*i (19) WA: Load Control Store State A. State WA is entered from State MO of a microinstruction which directs loading of microinstruction into FUSITT 11012. WA State Flag controls selection of addresses to CSADR Bus 20204 for writing into FUSITT 11012, and generates a write enable pulse to FUSITT 11012 to write microinstructions into FUSITT 11012.

(20) WB: Load Control Store State B. State WB is entered from State WA and is used to generate an appropriate timing interval for writing into FUSITT 11012. State WB also extends State Ml to 2 clock cycles to ensure a valid address input to FUSITT 11012 when a next microinstruction is to be read from FUSITT 11012.

Having described certain CS 10110 states, and S, operations which may be performed within those states, state sequences for certain CS 10110 operations will be described next below with aid of Figs. 244A to 244Z.

Fig. 244A to Fig. 244Z represent those state timing sequences necessary to indicate major features of CS 10110 state timing. All state timing shown in Figs. 244A to 244V assumes full pipelining of CS 10110 operations, for example pipelining of reads from and writes to MEM -344- 10112 by JP 10114. Pipelining is not assumed in Figs.

244W to 244Z. Referring to Figs. 244A to 244Z, these figures are drawn in the form of timing diagrams, with time increasing from left to right. Successive horizontally positioned "boxes" represents successive CS 10110 states during successive CS 10110 110 nano-second clock cycles. Vertically aligned "boxes" represent alternate CS 10110 states which may occur during a particular clock cycle. Horizontally extended dotted 10 lines connecting certain states represented in Fig. 244A to 244Z represent an indeterminate time interval which is an integral multiple of 110 nano-second CS 10110 clock S' cycles.

Referring to Fig. 244A to 244Z in sequence, State Timing Sequences shown therein represent: Fig. 244A; state timing for execution of a normal microinstruction with no Events occurring and no MEM 10112 references.

Fig. 244B; execution of a normal microinstruction, with no Events occurring, no MEM 10112 references, and a hold in State MO for one clock cycle.

Fig. 244C; a microinstruction requests an extension of State MO for one clock cycle, with no Events occurring and no MEM 10112 references.

Fig. 244D; a write to MEM 10112 from DESP 20210, for example from GRF 10354 or from OFFALU 20242. MEM 10112 port is available and MEM 10112 reference is made during first sequential occurrence of States MO and Ml.

-345- Fig. 244E; a write to MEM 10112 from DESP 20210 as described above. MEM 10112 port is unavailable for an indeterminate number of clock cycles. A MEM 10112 r-ference is made during first sequential occurrence of States MO and Ml.

Fig. 244F; writing of an EU 10122 result back into MEM 10112. MEM 10112 is available and a write operation is initiated during first sequential occurrence of States MO and Ml.

1 0 Fig. 244G; writing back of an EU 10122 result to MEM 10112 as described above. MEM 10112 port is -t unavailable for an undetermined number of clock cycles, or EU 10122 does not have a result ready to be written into MEM 10112. Write operation is initiated during first sequential occurrence of States MO and Ml.

Fig. 244H; a read of an EU 10122 result into FU 10120. EU 10122 result is not available for an undetermined number of clock cycles.

Fig. 2441; a read from MEM 10112 to OFFMUXR 23812, with no delays. The microinstruction following the microinstruction initiating a read from MEM 10112 does not reference OFFMUXR 23812.

Fig. 244J; a read from MEM 10112 to OFFMUXR 23812 with data from MEM 10112 being delayed by an indeterminate number of clock cycles. The next following microinstruction from that initiating the read from MEM 10112 does not reference OFFMUXR 23812.

(11) Fig. 244K; a read from MEM 10112 to O(FMUXR 23812. The next microinstruction following thE -346microinstruction initiating the read from MEM 10112 references OFFMUXR 23812.

(12) Fig. 244L; a read from MEM 10112 to GRF 10354.

The read to GRF 10354 is initiated by the first sequentially occurring States MO and Ml.

(13) Fig. 244M; a read from MEM 10112 to GRF 10354 and to OFFMUXR 23812. In this case, read operations may not be overlapped.

(14) Fig. 244N; JP 10114 honors an Event request and 10 initiates a corresponding Event Handler microinstruction Csequence, no MEM 10112 references occur.

(15) Fig. 2440; JP 10114 honors an Event request as stated above. MEM 10112 references are made during the first sequential occurrence of States MO and Ml and a MEM 10112 reference Event occurs. In case of an MEM 10112 reference event, State MA is entered from one clock cycle. This occurs only if a MEM 10112 reference is made and aborted.

(16) Fig. 244P; an Event occurs in a MEM 10112 S" 20 reference made during the first sequential occurrence of States MO and MI. The MEM 10112 reference does not result in a memory reference Event. CS 10110 remains in ~t State AB until the MEM 10112 reference is completed by return of data from MEM 10112.

(17) Fig. 244Q; a read of data from MEM 10112 or JPD SBus 10114 to EU 10122 Operand Queue. EU 10122 Operand Queue is not full.

(18) Fig. 244R; a read of MEM 10112 or JPD Bus 10142 data to EU 10122 Operand Queue. EU 10122 Operand Queue L Jstiaeo rercEe.C11rai i -347is full when the microinstruction initiating the read is issued.

(19) Fig. 244S; a request for a "nano-interrupt" to EU 10122 by FU 10120 with no Events occurring.

(20) Fig. 244T; FU 10120 submits a "nano-interrupt" request to EU 10122 and an EU 10122 State Overflow, described further in a following description, occurs. No other Events are recognized, as described in a following description of EVENT 20284.

S 10 (21) Fig. 244U; FU 10120 submits a "nano-interrupt" request to EU 10122. Another Event is recognized during State MO and an abort results. First abort state is j entered for the non-EU 10122 event All aborts 1 recognized in State MO are taken or acknowledged, before entrance into State MO. Therefore, on retry at State MO of the'original microinstruction entered from State MO, next abort recognized is for EU 10122 Stack Overflow r1 Event since EU 10122 Stack Overflow has higher priority.

(22) Fig. 244V; a load of a 27 bit microinstruction S 20 segment into FUSITT 11012.

In Figs. 244A to 244V, pipelining MEM 10112 reads and writes, and of JP 10114 operations, has been assumed. In 4. 4 Figs. 244W to 244Z, non-overlapping operation of JP 10114 *t tis assumed.

(23) Fig. 244W; a read of data from ME M 10112 to OFFMUXR 23812.

(24) Fig. 244X; a read of data from MEM 10112 to EU 10122 Operand Queue.

Fig. 244Y; a write of an EU 10122 result into MEM 10112.

(26) Fig. 244Z; a read of a 32 bit SIN word from MEM 10112 in response to a prefetch or conditional prefetch request.

Having described the general structure and operation of STATE 20294, and the operating states and operations of CS 10110, structure and operation of EVENT 20284 will be described next below.

b.b.b. Event- blogi c--20-2B4 -flg s---24 T7--24-7- S 10-2 An Event is a request for a change in sequence of execution of microinstructions which is generated by CS 10110 circuitry, rather than by currently executing r t microinstructions. Occurrence of an Event will result in provision of a microinstruction sequence, referred to as an Event Handler, by FUSITT 11012 which modifies CS 6 r 10110's operations in accordance with the needs of that oEvent. Event request signals may be generated by CS S, 10110 circuitry internal to JP 10114, that is from FU S 20 10120 or EU 10122 or CS 10110 circuitry external to JP 10114, for example from IOP 10116 or from MEM 10112.

Event request signals are provided as inputs to, EVENT 20284. As will be described further below, EVENT 20284 S masks Event Requests to determine which Events will be recognized during a particular CS 10110 Operating State, assigns priorities for servicing multiple Event Requests, and fabricates Handler addresses to FUSITT 11012 for microinstruction sequences for servicing requests. EVENT 20284 then provides those Handler microinstruction addresses to FUSITT 11012 through EVNTGT 24310, to a i 1 -349initiate execution of selected Event Handler microinstruction sequences.

Certain terms and expressions are used throughout the following description. The following paragraphs define these usages and provide examples illustrating these terms. An Event "makes a request" when a condition in CS 10110 hardware operation results in a Event Request signal being provided to EVENT 20284. As will be described further below, these Event Request signals are 1 0 provided to EVENT 20284 combinatorial logic which W determines the validity of those "requests".

L

An Event Request "is reognized" if it is not masked, ttvtt that is inhibited from being acted upon. Masking may be S. explicit, using masks generated by FUSITT 11012, or may be implicit, resulting from an improper CS 10110 State or invalid due to other considerations. That is, certain Events are recognized only during certain CS 10110 States even though those requests may be recognized during certain other states. Any number of requests, for 20 example up to 31, may be simultaneously recgized.

An Event Request is "honored" if it is the highest priority Event Request occurring. W'en a request is -honored, a corresponding address, of a corresponding t microinstruction sequence in FUSITT 11012, for its Handler microinstruction sequence is gated onto CSADR BUS 20204 by EVENT 20284. A request is honored when CS 10110 enters State AB. State AB gates the selected Event Handler microinstruction address on CSADR Bus 20284.

E -350- To summarize, a number of Events may request service by JP 10114. Of these Events, all, some, or none, may be recognized. Only one Event Request, the highest priority Event Request, will be honored when JP 10114 enters State AB. Microinstruction control of CS 10110 will then transfer to that Event's Handler microinstruction sequence. A necessary condition for entering State AB is that an Event Request has been made and.recognized.

A microinstruction sequence "completes", "is 10 completed", or reaches "completion" when CS 10110 exits State Ml while that microinstruction sequence is active.

k A microinstruction sequence may, as described above, be aborted in State MO an indefinite number of times before, if ever, reaching completion.

A MFM 10112 reference "completes", "is completed", or reaches "completion" when requested data is returned to th: specified destination, that is read from MEM 10112 to the requestor, or MEM 10112 accepts data to be written ,into MEM 10112.

"Trace Traps" are an inherent feature of microinatructions being executed. Trace Traps occur on every microinstruction of a given type (if not masked), Sfor example during a sequence of microinstructions to S perform a Name evaluate or resolve, and occur on each microinstruction of the sequence. In general, a Trace Trap Event must be serviced before execution of the next microinstruction. Trace Traps are distinct from Interrupts in that an Interrupt, described below, does not occur on execution of each microinstruction of a microinstruction sequence, but only on those L 'd 1 si L. ilr i ii^ -351microinstructions where certain other conditions must be considered.

"Interrupts" are the largest class of events in JP 10114. Occurrence of an Interrupt may not, in general, be predicted for a particular execution of a particular microinstruction in a particular instance. Interrupts may require service before execution of the next microinstruction, before execution of the current microinstruction can complete, or before beginning of the )10 next SIN. An Interrupt may be unrelated to execution of any microinstruction, and is serviced before beginning of the next microinstruction.

A "Machine Check" is an Event that JP 10114 may not S^ handle alone, or whose occurrence makes further actions by JP 10114 suspect. These events are captured in EVENT S20284 Registers and result in a request to DP 10118 to stop operation of JP 10114 for subsequent handling.

In summary, three major classes of Events in CS 10110 are Trace Traps, Interrupts, and Machine Checks. Each of S 20 these class of events will be described in further detail below, beginning with Trace Traps.

*4I4 The State of all possible Trace Trap Event Requests, Swhether requesting or not requesting, is loaded into EVENT 20284 Registers at completion of State Ml and at completion of State AB. That is, since Trap Requests are a function of the currently executing microinstruction, the State of a Trap Request will be loaded into EVENT 20284 Trace Trap Registers at end of State Ml of each currently executing microinstruction. Similarly, if any Trap Requests are recognized, State AB will be entered at I l -352the end of the first clock cycle of the next following State MO and their State loaded at end of the State AB.

Recognized, or unmasked, Trap Requests may be pushed onto RCWS 10358 as Pending Requests. Unrecognized, or masked, Trace Trap Requests may be pushed onto RCWS 10358 as Not Pending Requests and are subsequently disregarded. Subsequently, when a microinstruction sequence ends in a return to a calling microinstruction sequence, the Trace Trap Request bits in an RCWS 10358 10 may be used to generate Trace Trap Event Requests.

Upon exit from State AB, all Trace Trap Requests, except Micro-Break-Point and Microinstruction Trace Traps, described below, are loaded into corresponding EVENT 20284 Trace Trap Request Registers as not requesting. Micro-Break-Point and Microinstruction Trace Traps, are, in general, always latched as requesting at L t 1 completion of State AB. Trace Traps may be explicitly masked by a Trace Mode Mask, an Indivisibility Mode Mask, and by a Trace Enable input, all generated by FUSITT 11012 as described below. Micro-Break-Point Trap may also be masked by clearing a Trace Enable bit in a Trace Enable field of certain microinstructions containing Trace Traps. In general, masking is effective from State SMO of the microinstruction which generates the mask, through completion of a microinstruction which clears the mask Trace Traps generated by a microinstruction which clears a mask are taken so as to abort a following microinstruction during its MO State.

Referring to Fig. 245, CS 10110 state timing for a typical Trap Request, and generation of a i- -x -353microinstruction address to a corresponding Trace Trap Handler microinstruction sequence by EVENT 20284 is shown. Fig. 245 is drawn using the same conventions as described above with reference to Fig. 244A to 244Z. In Fig. 245, a microinstruction executing in States MO and Ml causes a Trace Trap Request but does not generate an MR (Memory Reference) Trailer State. Trace Trap Request to EVENT 20284 is signaled by Time A. This Trace Trap Request is latched into EVENT 20284 Trace Trap Event tie.

Registers, and an Abort Request is provided to STATE /L 20294. At Time B, FU 10120 enters States AB and AR. The i* t microinstruction address for a Handler microinstruction a Cs sequence of the highest priority Event present in EVENT 20284 is presented to FUSITT 11012 and execution of the addressed microinstruction sequence begins. At Time C, FU 10120 exits States AB and AR and enters State AB.

State AB will be exited at end of the next 110 nanosecond clock cycle. Address of the selected Event Handler microinstruction sequence will remain on CSADR Bus 20204 for duration of State AB. At Time D, a pointer into RCWS 10358, described in a following description, is ti 4. incremented, thereby effectively pushing the first microinstruction's return control word, that is the i microinstruction executing at first State MO, onto RCWS 10358. First microinstruction of the Trace Trap Event Handler microinstruction sequence is provided by FUSITT l 11012. Execution of Handler microinstruction sequence I, will begin at start of the third State MO of the state i t -354timing sequence shown in Fig. 245. EVENT 20284's Trace Trap Register for this event is now latched in nonrequesting state and will remain so until transition out of second State M1 shown in Fig. 245. At this time, EVENT 20284 Registers will latch new Trap Requests.

Finally, at Time E, Trace Trap Event Registers of EVENT 20284 are latched with new Trap Requests arising from execution of the microinstruction being executed in States MQ and M1 occurring between Times D and E. Traps due to the microinstruction that was executed in States MO and M1 before Time A, but were not serviced, are requested again when the previously pushed RCW described above is returned from RCWS 10358 upon return from the Trace Trap Event Handler microinstruction sequence initiated at Time D. All Trace Trap Requests which have been serviced are explicitly cleared in RCWS 10358 RCWs by their Event Handler microinstruction sequences to prevent recurrence of those Trap Requests. Since Trace Trap Event Requests arising from reads or writes to MEM 10112 will recur if those requests are repeated, EVENT 20284 generates memory repeat Interrupts after all tt aborted MEM 10112 read and write requests to insure that <these Traps will eventually be serviced. Event Handler i microinstruction sequences for these read and write Trace Trap Events explicitly disable serviced Trace Trap Event Requests by clearing bits in the logical descriptor of i! the aborted memory read and write requests.

Having described overall structure and operation of Trace Trap Events, certain specific Trace Trap Events will be described in greater detail below.. Trace Trap Events occurring in CS 10112 may include Name Trace 2I 1 *V t -355- Traps, SOP Trace Traps, Microinstruction Trace Traps, Micro-Break-Point Trace Traps, Logical Write Trace Traps, Logical Read Trace Traps, UID Read Trace TrapS, and UID Write Trace Traps. These Trace Traps will be described below in the order named.

A Name Trace Trap is requested upon every microinstruction sequence that contains an evaluate or resolve of a Name syllable. Name Trace Traps are provided by decoding certain microinstruction fields of S" i10 those microinstruction sequences. Name Trace Trap field is masked by either Trace Mask, Indivisibility Mask, or STrace Enable, as described above. All of these masks are tr, set and cleared by microinstruction control signals Sprovided during microinstruction sequences calling for resolves or evaluates of Name syllables.

A SOP Trace Trap may be requested whenever FU 10120 enters State FM (First Microinstruction of an SOP). SOP Trace Traps may be masked by Trace Mask, Indivisibility Mask, or Trace Trap Enable, again provided by microinstruction control outputs of FUSITT 11012. In general, the first microinstruction of such a microinstruction sequence interrupting such SOPs is not completed before a Trace Trap is taken.

Microinstruction Trace Traps may be requested upon completion of microinstructions which do not contain a Return Command, that is those microinstructions which do not return microinstruction control of CS 0110 to the calling microinstruction sequence. For microinstruction sequences containing Return Commands, state of microinstruction Trace Trap Request in a corresponding t T -356- RCW is used. Every microinstruction for which a Microinstruction Trace Trap is not masked is aborted during State MO of execution of that microinstruction.

Microinstruction Trace Traps may be masked by Trace Mask, Indivisibility Mask, or Trace Enable from FUSITT 11012.

A Micro-Break-Point Trap may be requested upon execution of microinstructions which do not contain Return Commands, but in which a Trace Enable bit in a microinstruction is asserted. A Micro-Break-Point Trap may be masked by Trace Mask, Indivisibility Mask, or f Trace Enable. In addition, a Trace Enable bit of a microinstruction field in these microinstruction sequences controls recognition of Micro-Break-Point f r- Traps. Micro-Break-Point Traps are thereby requested whenever a microinstruction Trace Trap is requested, but have additional enabling conditions expressed in the microinstructions. Since only recognized Traps are ^lo pushed onto RCWS 10358 in a RCW, a Microinstruction Trace Trap and a Micro-Break-Point Trap having different C 20 request states may be present in RCWS 10358 concurrently.

Logical Write Trace Traps may be requested when 4 enabled by a bit set in a logical descriptor during a t microinstruction sequence submitting a write request to I .MEM 10112 and using logical descriptors to do so.

Logical Write Trace Traps are recognized only if they occur during a state which will be immediately followed by State MR (Memory Reference Trailer). A Logical Write Trace Trap will result in the MEM 10112 write request -357being aborted. Logical Write Trace Traps may be masked by Trace Masks, Indivisibility Mask, or Trace Trap Enable. A further condition for recognition of a Logical Write Trace Trap is determined by the state of certain bits in a logical descriptor of the memory write request. Logical Write.Trace Traps are, in general, not pushed onto RCWS 10358 as part of a RCW since aborted MEM 10112 requests are re-generated so that Logical Write Trace Traps may be repeated.

10 Logical Read Trace Traps are similar in all respects to the Logical Write Trace Traps, but occur during MEM 10112 read requests. Generation of Logical Read Trace t Traps is controlled again in part by certain bits in tt logical descriptors of MEM 10112 read requests.

In certain implementations of CS 10110, UID Trace Traps may be requested when FU 10120 requests an MEM \10112 read operation based upon a UID address or pointer. UID Read Trace Traps are recognized if requested and there is, in general, no explicit masking of UID Read Trace Traps. Generation of UID Read Trace Traps is controlled by certain bits in MEM 10112 read request logical descriptors. UID Read Trace Trap c Requests result in the MEM 10112 read requests being aborted and CS 10110 entering State AB. Handler microinstruction sequences for UID Read Trace Traps will, in general, reset the trapped enable bit in the MEM 10112 read request logical descriptor before re-issuing the MEM 10112 read request.

ri "-ra cI-c -358- UID Write Trace Traps are similar to UID Read Trace Traps, and are controlled by bits in the logical descriptor in MEM 10112 write request based upon UID addresses or pointers.

Having described above structure and operation of Trace Trap Events, CS 10110 Interrupt Events will be described next below.

As previously described, Interruptsform the largest .class of CS 10110 Events. Interrupts may be regarded as S. 10 falling into one or more of several classes. First, Memory Reference Repeat Interrupts are those Interrupt r at Events associated, in general, with read and write requests to MEM 10112 in which a read or write request is S" submitted to MEM 10112, and an Interrupt Event results.

That Interrupt Event is handled, and the MEM 10112 request repeated. Second, Deferred Service Interrupts are those Interrupts wherein CS 10110 defers service of an Interrupt until entry to a new SIN. Fourth, Microinstruction Service Interrupts occur when a j| t currently executing microinstruction requires assistance of an Event Handler microinstruction sequence to be S t completed. Finally, Asynchronous Interrupt Events may I II occur at any time and must be serviced before CS 10110 may exit State MO of the next microinstruction. These Interrupt Events will be described next below in the order named.

A Memory Keference Repeat Interrupt is requested, for example, if a microinstruction executes a command, and a corresponding RCW read from RCWS 10358 indicates that a memory reference was aborted before entrance to the -359microinstruction sequence from which return was executed. This type of Interrupt Event occurs for all aborted memory references. If an event is honored, that is abort state is entered, for any event and there is a memory reference outstanding, not aborted, the memory reference completes before State AB is exited. No memory Repeat Interrupt Request will be written into the RCW written onto RCWS 10358. Conversely, if a memory reference is aborted, even if the event honored is not that event which aborted the memory reference, a Memory Repeat Interrupt Request will be written into a RCW t pushed onto a RCWS 10358.

There are two state timing sequences for execution of Memory Repeat Interrupts. In the first case, there are no MEM 10112 references in the mcroinstrUction executing a Return Command. In the second case, a microinstruction Si executing a Return Command executes a return and also makes a MEM 10112 reference. Referring to Fig. 246, a CS 10110 State Timing Diagram for the first case is shown.

Fig. 246 is drawn using the same conventions as used in Fig. 244 and 245. As described above, in the first case a microinstruction executing a Return Command is executed ^in States MO and Ml following Time D. An aborted MEM S10112 reference was made in States MO and Ml preceding Time A. An MEM 10112 Reference Abort Request is made upon CS 10110's entry into State MR following Time A.

Since a Memory Repeat Interrupt is requested only from a I RCW provided by RCWS 10358, a Memory Repeat Interrupt is .LjIIU lb-- -360indicated only if a microinstruction executes a Return Command resulting in RCWS 10358 providing such an RCW.

Therefore, a Memory Repeat Interrupt Request Register of EVENT 20284 is loaded with "not requesting" at this time. At Time B, CS 10110 enters State AB, State AR, and State MA. At this time, a Memory Reference Abort Request is asserted and written into an RCW when State AB is exited just before Time D. At Time D, CS 10110 exits 1 State AR and State MA. As just described, CS 10110 will 10 remain in State B until Time D. At Time D, Memory SReference Abort Request is written into RCWS 10358 as part of an RCW and, as described further below, various RCWS 10358 Stack Pointers are incremented to load that RCW into RCWS 10358. At this time, EVENT 20284's Interrupt Request Register receives "no request" as state of Memory Repeat Interrupt. First microinstruction of Memory Repeat Interrupt Handler microinstruction sequence is provided by FUSITT 11012. At Time E, the last t. microinstruction of the Memory Repeat Interrupt Handler microinstruction sequence is provided by FUSITT 11012 and a Return Command is decoded. RCWS 10358 Previous Stack Pointer, previously desribed, is selected to address RCWS ,,10358 to provide the previously written RCW as output to SEVENT 20284's Memory Repeat Interrupt Event Register. At Time F, EVENT 20284's Memory Repeat Interrupt Register is loaded from output of RCWS 10358 and RCWS 10358's Stack Register Pointers are decremented. At this time, Memory Repeat Interrupt Request is made and, as described below, -361is written into the current Return Control Word, whether honored or not. JP 10114 then repeats the aborted MEM 10112 reference.

In the second case, a State Timing Sequence wherein the microinstruction executing a return also makes a MEM 10112 reference, CS 10110 State Timing is identical up to Time F. At Time F, MEM 10112 Repeat request is not recognized and the state of Memory Repeat Interrupt written into the current'Return Control Word is "not 10 requesting" unless a current MEM 10112 reference is 4. aborted. The previous MEM 10112 Repeat Interrupt Request is disregarded as it is assumed that it is no longer "required. Thus, there are two ways to avoid, or cancel a Memory Repeat Interrupt Request. First, that portion of a RCW receiving a MEM 10112 Repeat Interrupt Request may be rewritten as "not requesting". Second, an aborted MEM 10112 reference may be made in the same microinstruction S' that returns from a Handler servicing the aborted MEM 10112 reference.

Certain CS 10110 Events result in aborting a MEM 10112 read or write references and may result in repeat of MEM 10112 references. These events may include: Logical read'and write Traps and, in certain implementations of CS 10110, UID read and write Traps, previously discussed; A PC 10234 miss; Detection of a Protection Violation by PC 10234; A Page Crossing in a MEM 10112 read or write request; w 1 -362- A Long Address Translation, that is an ATU 10228 miss requiring JP 10114 to evaluate a logical descriptor to provide a corresponding physical descriptor; Detection of a reset dirty bit flag from ATU 10228 upon a MEM 10112 write request as previously described; An FU 10122 stack overflow; 9 4n 4* 4) 44 4 4*4*r 44 4 44( 44i 4 44 (8) (9) An FU 10122 Illegal Dispatch; A Name Trace Trap event as previously described; A Store Back Exception, as will be described below; (11) EU 10122 Events resulting in aborting of a Store Back, that is a write request to MEM 10112 from EU 10122; (12) A read request to a non-accelerated Stack Frame, 15 that is a Stack Frame presently residing in MEM 10112 rather than accelerated to JP 10114 Stack Mechanisms; and, (13) Conditional Branches in SIN sequences resulting outsanding MEM 10112 read reference from PREF 20260; and, Of these Events, Logical Read and Write Traps, UID Read and Write Traps, and Name Trace Traps have been previously described. Other Events listed above will be described next below in further detail.

A PC 10234 Miss Interrupt may be requested upon a 25 logical MEM 10112 reference, that is when a logical descriptor is provided as input to ATU 10228 and a protection state is not encached in PC 10234. PC 10234 will, as previously described, indicate that a ii ^t a -363corresponding PC 10234 entry is not present by providing a Event Protection Violation (EVENTPVIOL) output to EVENT 20284. PC 10234 will concurrently assert an Abort output (ABORT) to force CS 10110 into State AB and thus abort that MEM 10112 reference.

A Page Crossing MEM 10112 Reference Interrupt is requested if a logical MEM 10112 reference, that is a logical descriptor, specifies an operand residing on two logical pages of MEM 10112. An output of ATU 10228 will S 10 abort such MEM 10112 references by asserting an Abort output (ABORT).

SA Protection Violation Interrupt is requested if a logical MEM 10112 reference does not possess proper "access rights, a mode violation, or if that reference appears to refer to an illegal portion of that object, an Su extent violation. Again, PC 10234 will indicate B* occurrence of a Protection Violation Event, which may be it disabled by a microinstruction control output of FUSITT 11012.

A Long Address Translation Event may be requested upon a logical MEM 10112 reference for which ATU 10228 does not have an encached entry. ATU 10228 will abort that MEM 10112 reference by asserting outputs ABORT and Long Address Translation Event (EVENTLAT).

A Dirty Bit Reset Event Interrupt may be requested when JP 10114 attempts to write to an MEM 10112 page having an encached entry in ATO 10228 whose dirty bit is not set. ATU 10228 will abort that MEM 10112 write 1 e -364request by asserting outputs ABORT and Write Long Address Translation Event (EVENTWLAT).

An FU 10120 User Stack Overflow Event may be requested if the distance between a Current Frame Pointer and a Bottom Frame Pointer, previously described with reference to CS 10110 Stack Mechanisms, is greater than a given value. As previously described, in CS 10110 this value is eight. A User Stack Overflow Event will continue to be requested until either Current Frame Pointer or Bottom Frame Pointer changes value so that the difference limit defined above is no longer violated. A User Stack Overflow Event may be masked by a Trace Mask, an Indivisibility Mask, or by enable outputs of a microinstruction from FUSITT 11012. A Handler microinstruction sequence for User Stack Overflow Events must be executed with one or more of these masks set to :prevent recursion of these events. CS 10110 is defined 1' to be running on Monitor Stack (MOS) 10370 when User Stack Overflow Events are masked. User Stack Overflow Events are nob loaded into any of EVENT 20284's Event Registers, nor are these events written into a RCW to be written onto RCWS 10358.

Illegal EU 10122 Dispatch Events are requested by SSUSDT 20266 if FU 10120 attempts to dispatch, or provide 25 an initial microinstruction sequence address, to EU 10122 to a EUSITT address which is not accessable to a user's program. Illegal EU 10122 Dispatch Events are, in general, not maskod. Illegal EU 10122 Dispatch Event Requests are cleared upon CS 10110 exits from State AB.

-365- The Handler microinstruction sequence for Illegal EU 10122 Dispatch Events should, in general, reset Illegal EU 10122 Dispatch Event entries in RCWs to prevent recursion of these events.

EU 10122 will indicate a Store Back Exception Event if any one of a number of exceptional conditions arise during arithmetic operations. These events are recognized when CS 10110 enters State SB and are ignored except during Store Back to MEM 10112 of EU 10122 results. These Events may be disabled by microinstruction output of FUSITT 11012 but are, in general, not masked. Store Back Exception Events may be •written into RCWs, to be stored in RCWS 10358, and are cleared upon CS 10110's exit from State AB. Again, a Store Back Exception Event Handler microinstruction sequence should reset Store Back Exception Events written into RCWs to prevent recursion of these events.

As described above, the next major class of Interrupt Events are Deferred Service Interrupts. CS 10110 defers service of Deferred Service Interrupts until entry of a new SOP Deferred Service Interrupts which have been recognized will be serviced before completion of execution of the first microinstruction of that new SOP.

Deferred Service Interrups include Nonfatal MEM 10112 Errors, Interval Timer Overflows, and Interrupts from IOS 10116. These Interrupts will be described below, in the order named.

A Nonfatal MEM 10112 Interrupt is signaled by MEM 10112 upon occurrence of a correctable (single bit) MEM 10112 error. Nonfatal Memory Error Interrupts are u l -366recognized only during State MO of the first microinstruction of an SOP. MEM 10112 will continue to assert Nonfatal Memory Error Interrupt until JP 10114 issues an acknowledgement to read MEM 10112's Error Log.

An Interval Timer Overflow Interrupt is indicated by TIMERS 20296 when, as described below, an Interval Timer increments to zero, thus indicating lapse of an allowed time limit for execution of an operation. Interval Timer Overflow Interrupts are recognized during State MO of the 10 first microinstruction of a SOP. TIMERS 20296 will continue to request such interrupts until cleared by a microinstruction output of FUSITT 11012.

S" IOS 10116 will indicate an IOS 10116 Interrupt to indicate that an inter-processor message from IOS 10116 15 to JP 10114 is pending. IOS 10116 will continue to t cC assert an IOS 10116 Interrupt Request, which is stored in a register, until cleared by a microinstruction control output of FUSITT 11012. IOS 10116 Interrupts are S, recognized during State MO of the first microinstruction of an SOP.

The next major class of CS 10110 events are

S

t Interrupts due to the requirement by microinstruction "l sequences to be serviced in order to complete execution.

These Interrupts must be serviced before a microinstruction sequence may be completed.

Microinstruction Service Interrupts include Illegal SOP Events, Microinstructions Not Present in FUSITT 11012 Events, an attempted parse of a hung INSTB 20262, underflow of an FU 10120 Stack, an NC 10226 Cache Miss, or an EU 10122 Stack Overflow. Each of these events will r v w~l -367be described below, in the order named.

An Illegal SOP Event is indicated by FUSDT 11010 to indicate that a current SOP Code is a Long Code, that is greater than eight bits, while the current dialect (S- Language) expects only Short Operation Codes, that is eight bit SOPs. An Illegal SOP Interrupt is not detected for unimplemented SOPs within the proper code length range. Illegal SOP Events are, in general, not masked.

FUSDT 11010 continues to indicate an Illegal SOP Event 10 until a new SOP is loaded into OPCODEREG 20268. Illegal S. SOP Events are recognized during the first microinstruction of an SOP, that is during State FM.

.Should a Handler microinstruction sequence for a higher priority event change contents of OPCODEREG 20268, a previous Illegal SOP Event will be indicated again when 0 the aborted SOP is retried.

Absence of a Microinstruction in FUSITT 11012 is indicated by FUSITT 11012 asserting a Control Store S' Address Invalid (CSADVALID). This FUSITT 11012 output indicates that that particular microinstruction address points outside of FUSITT 11012's address space. Output of FUSITT 11012 in such event is not determined and parity checking, described below, of microinstruction Si output is inhibited. The Handler microinstruction sequence for these Events will load FUSITT 11012 address zero with the required microinstruction from MEM 10112, as previously described, and return to the original microinstruction sequence.

An attempted parse of a hung INSTB 20262 is indicated by INSTBWC 24110 when a parse operation is attempted, LI r si. -368- INSTB 20262 is empty, and PREF 20260 is not currently requesting SINs from MEM 10112. In general, these Events are not masked. If a higher priority Event is serviced, these Events are indicated again when the aborted microinstruction is retried if the original conditions still apply.

An FU 10120 Stack Underflow Event is requested when a current microinstruction references a Previous Stack Frame which is not in an accelerated stack, that is, the 10 Current Stack Pointer equals Bottom Stack Pointer. FU 10120 Underflow Events are, in general, not masked and are requested again on a retry if the microinstruction is aborted and this event has not been serviced.

An NC 10226 Miss Interrupt occurs on a MEM 10112 read or write operation when a load or read of NC 10226 is S' attempted and there is no valid NC 10226 block corresponding to that Name syllable. An NC 10226 Miss Event does not result in a request for a Name evaluate or V resolve. In general, these Events are not masked and result in a request being issued again if the microinstruction resulting in that Event is retried and has not been serviced.

t< An EU 10122 Stack Overflow Event is requested from EU 10122 to indicate that EU 10122 is currently already servicing at least one level of Interrupt an FU 10122 is requesting another. As will be described in a following description of EU 10122, EU 10122 contains a one level" i deep stack for handling of Interrupts. EU 10122 Stack -369- Overflow Events are enabled during State NW. All previously pending events will have been serviced before EU 10122 Stack Overflcw Event requests are recognized.

These Events will be serviced immediately upon entry into a following State MO, being the highest priority interrupt event. EU 10122 Stack Overflow Events may, in general, not be masked and once recognized are the next honored event.

"te" Finally, the third major class of CS 10110 Interrupt 10 Events are Asychronous Events. Asynchronous Events must, in general, be serviced before exiting State MO of a t.ta' microinstruction after they are recognized. Asychronous a a Events include Fatal Memory Error Events, AC Power Failure Events, Egg Timer Overflow Events, and EU 10122 Stack Underflow Events. CS 10110 Egg Timer is a part of Ci TIMERS 20296 and will be discussed as part of TIMERS 20296. These events will be described below, in the order referred to.

tt Fatal MEM 10112 Error Events are requested by MEM 10112 by assertion of control signal output PMODI, previously described, when last data read from MEM 10112 •contains a noncorrectable error. Fatal MEM 10112 Error t .Events are recognized on first State MO after occurrence. Fatal MEM 10112 Error Events are stored in an EVENT 20284 Event Register and are cleared upon entry into its service microinstruction sequence. In general, Fatal MEM 10112 Error Events may not be masked.

-370- AC Power Failure Events are indicated by DP 10118 by assertion of output signal ACFAAIL when DP 10118 detects a failure of power to CS 10110. Recognition of AC Power lailure Events is disabled upon entry to AC Power Failure Event Handler microinstruction sequence. No further AC Power Failure Events will be recognized until DP 10118 reinitiates JP 10114 operation.

As will be described further below, FUCTL 20214's Egg Timer is a part of TIMERS 20296. Egg Timer Overflow 10 Events are indicated by TIMERS 20296 whenever TIMERS 20296's Egg Timer indicates overflow of Egg Timer Counter. Egg Timer Overflow Events may be masked as I' described in a following description.

Finally, EU 10122 Stack Underflow Events are signaled by EU 10122 when directed to read a word from EU 10122 Stack Mechanism and there is no accelerated stack frame present. EU 10122 will continue to assert this Event Interrupt until acknowledged by JP 10114 by initiation of a n a Handler microinstruction sequence.

The above descriptions of CS 10110 events have stated that recognition of certain of those Events may be masked, that is inhibited to allow recognition of other p" Events having higher priority. Certain of these masking operations were briefly described in the above descriptions and will be described in further detail next below. In general, recognition of Events may be masked in five ways, four of which are properly designated as masks. These four masks are generated by microinstruction control from FUSITT 11012 and include Asychronous Masks for, in general, Asychronous Events.

4 i -371- Monitor Masks are utilized for those CS 10110 operations being performed on Monitor Stack (MOS) 10370, as previously described with reference to CS 10110 Stack Mechanisms. Trace Mask is utilized with reference to Trace Trap Events. Indivisible Mask is generated or provided by FUSITT 11012 as an integral or indivisible part of certain microinstructions and allow recognition of certain selected events during certain single microinstructions. Certain other Events, for example 10 Logical Read and Write Traps and UID Read and Write Traps, are recognized or masked by flag bits in logical j f, tdescriptors associated with those operations. Finally, certain microinstructions result in FUSITT 11012 providing microinstruction control outputs enabling or inhibiting recognition of certain events, but differ from Indivisible Masks in not being associated with single particular microinstructions.

Referring to Fig. 247, the relative priority level Stand applicable masks of certain CS 10110 Events are depicted therein in three vertical columns. Information regarding priority and masking of particular Events is t I 9 |shown in horizontal entries, each comprising an entry in t t each of these three vertical columns. Left hand column, titled Priority Level, states relative priority of each Event entry. Second column, titled EVENT, specifies which Event is referred to in that table entry. A particular Event will yield priority to all higher priority Events and will take presidence over all lower priority Events. Fig. 247's third column, titled Masked By, specifies for each entry which masks may be used to L, w w w i -372mask the corresponding Event. A indicates use of Asychronous Masks, M use of Monitor Mask, T use of Trace Trap Mask, and I represents that Indivisible Mask may be used. DES indicates that an Event is enabled or masked by flag bits of logical descriptors, while MCWD indicates that a particular Event may be masked by microinstruction control signal outputs provided by FUSITT 11012. NONE indicates that a particular Event may, in general, not be too* masked.

St, 10 The final major class of CS 10110 event was described ri above as Machine Check Events. In general, if any of r these Events are detected by logic gating in EVENT 20284, EVENT 20284 will provide a Check Machine signal to DP 10118. DP 10118 will then stop operation of JP 10114 and Machine Check Event Handler microinstruction sequences will be initiated. Among these Machine Check Events are Swherein FU 10120 is attemting to store back an EU 10122 result to MEM 10112 and EU 10122 signals a parity error V in EU 10122's Control Store. These events are stored in EVENT 20284 Event Registers and recogized when FU 10120 enters State AB. EU 10122 will have previously ceased t' operation until a corrective microinstruction sequence may be initiated. The same Event will occur if FU 10120 attempts to use an EU 10122 arithmetic operation result or test operation result having a parity error in EU 10122's Control Store. Shouid MOS 10370 overflow or underflow, this event will be detected, FU 10120 operations stopped, and corrective microinstruction 1 -373sequences initiated. MOS 10370 overflow or underflow occurs whenever a previous MOS 10370 Stack Frame is referenced, whenever MOS 10370 Stack Pointer equals MOS 10370 Bottom Stack Pointer, or the difference between MOS 10370 Current and Bottom Stack Pointers is greater than sixteen. Underflows result in a transfer of operation to MIS 10368, while overflows are handled by DP 10118.

Finally, a Machine Check Event will be requested when a parity error is detected in a microinstruction currently ti V 10 being provided by FUSITT 11012 during State MO of that microinstruction.

Having described general operation of EVENT 20284, S, the structure and operation of EVENT 20284 will be described briefly next below.

Referring to Fig. 248, a partial block diagram of EVENT 20284 is shown. EVENT 20284 includes Event Detector (EDET) 24810, Event Mask and Register Circuitry (EMR) 24812, and Event Handler Selection Logic (EHS) E24814. EDET 24810 is comprised of random logic gating and, as previously described, receives inputs representing event conditions from other portions of CS 10110's circuitry. EDET 24810 detects occurrences of CS W 10110 operating conditions indicating that Events have occurred and provides outputs to EMR 24812 indicating what Events are requested.

EMR 24812 includes a set of registers, for example SN74S194s, comprising EVENT 20284's Event Registers.

These registers are enabled by mask inputs, described momentarily, to enable masking of those Events which are latched in EVENT 20284's Event Registers. Certain I. .n, -374- Events, as previously described, are not latched and logic gating having mask enable inputs is provided to enable masking of those events which are not latched.

EMR 24812 mask inputs are Asychronous, Monitor, Trace Trap, and Indivisible Masks, respectively AMSK, MMSK, TMSK, and ISMK, provided from FUSITT 11012. Mask inputs derived from FUSITT 11012 microinstruction outputs (mWRD) are provided from microinstruction control outputs of FUSITT 11012. EMR 24812 provides outputs representing 10 mask and unmask events which have been requested to EHS 1^ 24814.

EHS 24814 is comprised of loic gating detecting which of EHS 24814's unmasked Event Requests is of highest priority. EHS 24814 selects the highest priority unmasked Event Request input and provides a corresponding Event Handler microinstruction address to EVNTGT 24310 through ADRA Bus 24322. These address outputs of EHS 24814 are five bit addresses selecting the initial microinstruction of the Event Handler microinstruction sequence of the current highest priority unmasked Event.

As previously described with reference to NASMUX 24312, certain inputs of ENTGT 24310 are hard-wired to provide a full fifteen bit address output from EVNTGT 24310. EVENT 20284 also provides, from EHS 24814, an Event Enable Select (EES) output to SITTNAS 20286 to enable EVNTGT 24310 to provide microinstruction addresses to CSADR Bus 20204 when EVENT 20284 must provide a microinstruction address for handling of a current Event.

-i -375- Having described the structure and operation of FUCTL 20214's circuitry providing microinstruction addresses to FUSITT 11012, FUSITT 11012 will be described next below.

c.c.c. Fetch--Unt--S- -nterprete-Table---fO-I0-2 (Pfia-.--2494 Referring to Fig. 249, a partial block diagram of FUSITT 11012 is shown. Address (ADR) and Data (DATA) inputs of Micro-Instruction Control Store (mCS) 24910 are 0 10 connected, respectively, from CSADR Bus 20204 through Address Driver (ADRDRV) 24912 and from JPD Bus 10142 through Data Driver (DDRV) 24194. mCS 24910 comprises a r. memory for storing sequences of microinstructions currently being utilized by CS 10110. mCS 24910 is an 8K (8192) word by 80 bit wide memory. That is, mCS 24910 may contain, for example, up to, 8192 80 bit wide microinstructions. Microinstructions to be written into mCS 24910 are provided, as previously described, to mCS 24910 DATA input from JPD Bus 10142 through DDRV 24914.

Addresses of microinstructions to be written into or read a ,from mCS 24910 are provided to mCS 24910 ADR input from 1., t i CSADR Bus 20204 through ADRDRV 24912. ADRDRV 24912 and DDRV 24914 are buffer drivers comprised, for example, of SN74S240s and SN74S244s.

Also connected from output of ADRDRV 24912 is input of Nonpresent Micro-Instruction Logic (NPmIS) 24916.

NPmIS 24916 is comprised of logic gating monitoring read addresses provided to mCS 24910. When a microinstruction p t -376- 9 9* n *r It I ItL $9 4* t 4i 9 *9 9 read address present on CSADR Bus 20204 refers to an address location not within mCS 24910's address space, that is of a non-present microinstruction, NPmIS 24916 generates an Event Request output indicating this occurrence. As previously described FUCTL 20214 will then call, and execute, microinstructions so addressed from MEM 10112.

As indicated in Fig. 249, mCS 24910 provides three sets of outputs. These outputs are Direct Output (DO), Direct Decoded Output (DDO), and Buffered Decode Output (BDO). In general, control information within a particular microstruction word is used on next clock cycle after the address of that particular microinstruction word has been provided to mCS 24910 ADR 15 input. That is, during a first clock cycle a microinstruction's address is provided to mCS 24910 ADR input. That selected microinstruction appears upon mCS 24910's DO, DDO, BDO outputs during that clock cycle and are used, after decoding, during next clock cycle.

Outputs DO, DDO, BDO differ in delay time before decoded microinstruction outputs are available for use.

mCS 24910 DO output provides certain bits of microinstruction words directly to particular destinations, or users, through Direct Output Buffer (DOB) 24918. These microinstructions bits are latched and decoded at their destinations as required. DOB 24918 may be comprised, for example, of SN74S04s.

mCS 24910's DDO output provides decoded microinstruction control outputs for functions requiring the presence of fully decoded control signals at the 999* 9.

4 4: 4 i44 -377start of the clock cycle in which those decoded control signals are utilized. As shown in Fig. 249, mCS 24910's DDO output is connected to input of Direct Decode Logic (DDL) 24920. DDL 24920 is comprised of logic gating for decoding certain microinstruction word bits during same clock cycle in which those bits are provided by mCS 24910's DDO. These microinstruction bits are provided, as described above, during the same clock cycle in which a corresponding address is provided to mCS 24910's ADR 10 input During this clock cycle, DDL 24920 decodes mCS S24910's DDO microinstruction bits to provide fully decoded outputa by end of this clock cycle. Outputs of DDL 24920 are connected to inputs of Direct Decode e Register (DDR) 24922. DDR 24922 is a register comprised, for example, of SN74S374s. DDL 24920's fully decoded outputs are loaded into'DDR 24922 at the end of the clock S- cycle during which, as just described, an address is Sprovided to mCS 24910's ADR input and mCS 24910's b, corresponding DDO output is decoded by DDL 24920. Fully decoded microinstruction control outputs corresponding to mCS 24910's DDO outputs are thereby available at start of r' the second clock cycle. Microinstruction control outputs of DDR 24922 are thereby available to FU 10120 at start of the second clock cycle for those FU 10120 operations requiring immediate, that is undelayed, microinstruction control signal outputs from FUSITT 11012.

Finally, mCS 24910's BDO is provided for those FU 10120 operations not requiring microinstruction control signals imediately at the start of the second clock cycle. As shown in Fig. 249, mCS 24910's BDO is i 1 1 ii 1 i -378connected to inputs of Buffered Decode Register (BDR) 24924. Microinstruction word output bits from mCS 24910's BDO are provided to inputs of BDR 24924 during the clock cycle in which a corresponding address is provided to mCS 24910's ADR input. mCS 24910's BDO outputs are loaded into BDR 24924 at end of this clock cycle. BDR 24924's outputs are connected to inputs of Buffered Decode Logic (BDL) 24926. BDL 24926 is comprised of logic gating for decoding outputs of BDR S* 10 24924. BDL 24926 thereby provides decoded microinstruction control outputs to FU 10120 at some I t delayed time after start of the second clock cycle.

SMicroinstruction control outputs from BDL 24926 are i thereby delayed in time from the appearance of microinstruction control outputs of DDR 24922 but, as BDR 24924 stores microinstruction word bits rather than decoded microinstruction word bits, BDR 24924 is required to store proportionately fewer bits than DDR 24922.

S Finally, as shown in Fig. 249 outputs of DDR 24922 and BDR 24924, are connected to inputs of Microinstruction Word Parity Checker (mWPC) 24928a mWPC 24928 is comprised of logic gating for checking parity of outputs of DDR 24922 and BDR 24924. A failure in parity of either output of DDR 24922 and BDR 24924 indicates a possible error in microinstruction output from mCS 24910. When such an error is detected by mWPC 24928, mWPC 24928 generates a corresponding Microinstruction Word Parity Error (mWPE). -379d.d. coxntro Associated with SR's 10362, the stack mechanism area of GRF 10354, are two CS 10110 control structures primarily associated with operation of CS 10110's internal mechanisms. A first of these referred to as Machine Control Block, describes current execution environment of JP 10114 microprograms, that is, JP 10114 microinstruction sequences. Machine Control Block is 10 comprised of two information words residing in MCWl 20290 S, and MCWO 20292. These Machine Control Words contain all control state information necessary to execute JP 10114's 1 current microprogram. Second control structure is a portion of RCWS 10358, which as previously described parallels the structure of SR's 10362. Each register frame on MIS 10368 or MOS 10370 has, with exception of Top (Current) Register Frame, associated with it a Return Control Word (RCW) residing in RCWS 10358. RCWs are created when MIS 10362 or MOS 10370 register frames are pushed, that is moved onto MIS 10368 or MOS 10370 due to creation of a new Current Register Frame. A current RCW does not exist in a present embodiment of CS 10110.

RCWS 10358 will be described first next below, followed by Machine Control Block.

a.a.a. Return-Contro+-Word Referring to Fig. 251, a diagramic representation of a RCWS 10358 RCW is shown. As previously described, RCWS 10358 RCWs contain information necessary to reinitiate or continue execution of a microinstruction sequence if i; 41 -380execution of that sequence has been discontinued.

Execution of a microinstruction sequence may be discontinued due to a requirement to service a CS 10110 Event, as described above, or if that microinstruction sequence has called for execution of another microinstruction sequence, as in a Branch or Case Operation.

As shown in Fig. 251, each RCW may contain, for to example, 32 bits of information. RCW Bits 16 to 31 10 inclusive are primarily concerned with storing current Si. microinstruction address of microinstruction sequeres which have been discontinued, as described above. Bits L/ 17 to 31 inclusive contain microinstruction sequence return address. Return address is, as previously described, address of the microinstruction currently S. being executed of a microinstruction sequence whose execution has been discontinued. When JP 10114 returns from servicing of an Event or execution of a called I microinstruction sequence, return address is provided from RCWS 10358 to SITTNAS 20286 and through CSADR Bus 20204 to FUSITT 11012 as next microinstruction address to resume execution of that microinstruction sequence. Bit 16 of an RCW contains a state bit indicating whether the particular microinstruction referred to by return address field is the first microistruction of a particular SOP.

That is, Bit 16 of an RCW stores CS 10110 State FM.

Bits 8 to 15 inclusive of an RCW contain information pertaining to current condition code of JP 10114 and to pending Interrupt Requests. In particular, Bit 8 contains a condition code bit which, as previously iJ -381described indicates whether a particular test condition has been met. RCW Bit 8 is thereby, as previously described, a means by which JP 10114 may pass results of a particular test from one microinstruction sequence to another Bits 9 to 15 inclusive of an RCW contain information regarding currently pending Interrupts.

These Interrupts have been previously discussed, in general, with reference to EVENT 20284. In particular, t 1..t RCW Bit 9 contains pending state of Illegal EU 10122 I 1 0 Dispatch Interrupt Requests; RCW Bit 10 contains pending state of Name Trace Trap Request; RCW Bit 11, contains .pending state of Store Back Interrupt Request; RCW Bit 12 contains pending state of Memory Repeat Interrupt Request; RCW Bit 13 contains pending state of SOP Trace Trap Request; RCW Bit 14 contains pending state of Microtrace Trap Request; and, RCW Bit 15 contains pending state of Micro-Break Point Trap Request. Interrupt Handling microinstruction sequence which require use of f. CS 10110 mechanisms containing information regarding pending Interrupts must, in general, save and store that information. This save and restore operation is accomplished by use of Bits 9 to 15 of RCWS 10358's RCWs. Upon entry to an Interrupt Handling microinstruction sequence, these bit flags are set to indicate Interrupts which were outstanding at time of entry to that microinstruction sequence. Because these bits are used to initiate Interrupt Request upon returns, pending Interrupts may be cancelled by resetting appropriate bits of Bits 9 to 15 upon return. This capability may be used to implement Microinstruction Trace Traps, previously described.

-382- As indicated in Fig. 251, RCW Bits 0 to 7 are not utilized in a present embodiment of CS 10110. RCW bits 0 to 7 are not implemented in a present embodiment of CS 10110 but are reserved for future use.

As previously described, RCWs may be writtten into or read from RCWS 10358 from JPD Bus 10142. This allows contents of RCWS 10358 to be initially written as desired, or read from RCWS 10358 to MEM.10112 and subsequently restored as required for swapping of o10 processes in CS 10110.

4 b.b.b. Mach-ine-e Con-ro--B-l-ock t As described above, FUCTL 2Q214's Machine Control c l Block is comprised of a Machine Control Word 1 (MCW1) and a Machine Control Word 0 (MCWO) MCW1 and MCWO reside, respectively, in Registers MCW1 20290 and MCWO 20292.

SMCW1 and MCWO described the current execution environment of FUCTL 20214's current microprogram, that is the •t microinstruction sequence currently being executed by JP 10114.

Referring to Fig. 252, diagramic representations of MCWO and MCW1 are shown. As indicated therein, MCWO and MCW1 may each contain, for example, 32 bits of information regarding current microprogram execution environment.

Referring to MCWO, MCWO includes 6 execution environment subfields. Bits 0 to 3 inclusive contain a Top Of Stack Counter (TOSCNT) subfield which is a pointer -383to Current Frame of accelerated Microstack (MIS) 10368.

TOSCNT field is initially set to point to Frame 1 of MIS 10368. Bits 4 to 7 inclusive comprise a Top of Stack -1 Counter (TOS-1CT) subfield which is a pointer to Previous Frame of accelerated MIS 10368, that is to the MIS 10368 frame proceeding that pointed by TOSCNT subfield. TOS- 1CNT subfield is initially set to Frame 0 of MIS 10368.

Bits 8 to 11 inclusive comprise a Bottom of Stack Counter (BOSCNT) subfield which is a pointer to Bottom Frame of accelerated MIS 10368. BOSCNT subfield is initially set to point to Frame 1 of MIS 10368. TOSCNT, TOS-1CNT, and BOSCNT subfields of MCWO may be read, written, incremented and decremented under microprogram control as frames are transferred between MIS 10368 and a SS 10336.

Bits 17 to 23 inclusive and Bits 24 to 31 inclusive of MCWO comprise, respectively, Page Number Register (PNREG) and Repeat Counter (REPCTR) subfields which, together, comprise a microinstruction address pointing to a microinstruction currently being written into FUSITT 11012.

Bits 12 to 15 inclusive of MCWO comprise an Egg Timer (EGGT) subfield which will be described further below with respect to TIMERS 20296. Bit 16 of MCWO is not utilized in a present embodiment of CS 10110 Referring to MCW1, MCW1 is comprised of four subfields. Of the 32 bits comprising MCW1, Bits 0 to inclusive and Bits 24 and 25 are not utilized in a present embodiment of CS 10110. Bit 16 is comprised of a i i, i 1 -384- Condition Code (CC) subfield indicating results of certain test conditions in JP 10114. As previously described CC subfield is automatically saved and restored in RCWS 10358 RCW's.

Bits 17 to 19 inclusive of RCW1 comprise an Interrupt Mask (IM) subfield. The three bits of IM subfield are utilized to indicate a hierarchy of non-interruptible JP 10114 microinstruction control operating states. That is, a three bit code stored therein indicates relative 10 power to interrupt between three otherwise non- Sinterruptible JP 10114 operating states. Bits 20 to 23 ,ttt inclusive comprise an Interrupt Request (IR) subfield St, which indicate Interrupt Request. These Interrupt Requests may include, for example, Egg Timer Overflow, Interval Timer Overflow, or Non-Fatal Memory Error, as SC have been previously described. Finally, Bits 26 to 31 Sinclusive comprise a Trace Trap Enable (TTR) subfield indicating which Trace Trap Events, previously described, are currently enabled. These enables may include Name Trace Enable, Logical Retrace Enable, Logical Write Trace Enable, SOP Trace Enable, Microinstruction Enable, and 4 Microinstruction Break Point Enable.

MCWO and MCW1 has been described above as if residing in registers having individual, discrete existence, that is MCW1 20290 and MCWO 20292. In a present embodiment of CS 10110, MCW1 20290 and MCWO 20292 do not exist as a unified, discrete register structure but are instead comprised of individual registers having physicl i existence in other portions of FUCTL 20214. MCW1 20290 and MCWO 20292, and MCW1 and MCWO, have been so described i L l L Ll l l -385to more distinctly represent the structure of information contained therein. In addition, this approach has been utilized to illustrate the manner by which current JP 10114 execution state may be controlled and monitored through JPD Bus 10142. As indicated in Fig. 202, MCW1 20290 and MCWO 20292 have outputs connected to JPD Bus 10142, thus allowing current execution state of JP 10114 to be read out of FUCTL 20214. Individual bits or subfields of MCWO and MCW1 may, as previously descibed, 10 be written by microinstruction control provided by FUSITT 11012. In a present physical embodiment of CS 10110, those registers of MCWO 20292 containing subfields TOSCNT, TOS-1CNT, and BOSCNT reside in RAG 20288. Those portions of MCWO 20292 containing subfield EGGT reside in TIMERS 20296. MCWO 20292 registers contain PNREG and Sr. REPCTR subfields are physically comprised of REPCTR 20280 and PNREG 20282. In MCW1 20290, CC subfield exists as output of FUCTL 20214 test circuits. Those MCW1 20290 V* registers containing IM, IR, and TTE subfields reside within EVENT 20284.

Having described FUCTL 20214 structure and operation as regards RCWS 10358, MCW1 20290 and MCWO 20292, FUCTL 20214, RAG 20288 will be described next below.

c.c.c. Register--Addess Generator--20-2-28-4 253 Referring to Fig. 253, a partial block diagram of RAG 20228, together with diagramic representation of GRF x

(L

-386- 10354, BIAS 20246 and RCWS 10358, is shown. As previously described, JP 10114 register and stack mechanisms include General Register File (GRF) 10354.

BIAS 20246, and RCWS 10358. GRF 10354 is, in a present embodiment of CS 10110, a 256 word by 92 bit wide array of registers. GRF 10354 is divided horizontally to provide Global Registers (GRs) 10360 and Stack Registers (SRs) 10362, each of which contains 128.of GRF 10354's 256 registers. GRF 10354, that is both GRs 10360 and SRs 10362, is divided vertically into three vertical sections S. designated as AONGRF 20232, OFFGRF 20234, and LENGRF 20236. AONGRF 20232, OFFGRF 20234, and LENGRF 20236 are, respectively, 28 bits, 32 bits, and 32 bits wide. GRs t 10360 is utilized as an array of 128 individual registers, each register containing one 92 bit word. SRs r 10362 is structured and utilized as an array of 16 S' register frames wherein each frame contains eight registers and each register contains one 92 bit wide t word. Eight of SR 10362's frames are utilized as Microstack (MIS) 10362 and the remaining eight of SR 10362's frames are utilized as Monitor Stack (MOS) 10370. For addressing purposes only, as described further below, GRs 10360 is regarded as being structured in the same manner as SRs 10362, that is as 16 frames of eight registers each.

BIAS 20246, as previously describe, is a register array within BIAS 20246. BIAS 20246 contains 128 six bit wide registers, or words, and operates in parallel with and is addressed in parallel with SR 10362 portion of GRF 10354. RCWS 10358 is, as previously described, an array r -387of 16 registers, or words, wherein each register contains one 32 bit RCW. RCWS 10358 is structured and operates in parallel with SRs 10362 with each RCWS 10358 register corresponding to a SR 10362 frame of eight registers. As described below, RCWS 10358 is addressed in parallel with SR 10362's frames.

Source and Destination Register Addresses (SDAR) for selecting a GRF 10354 register to be, respectively, read from or written to are provided by RAG 20288. As 10 described above BIAS 20246 operates and is addressed in parallel with SR 10362 portion of GRF 10354, that is parallel with SRs 10362. BIAS 20246 registers are thereby connected to and in parallel with address inputs of SRs 10362 and are addressed concurrently with GRs 10360. Registers RCWS 10358 also operate and are addressed in parallel with SRs 10362. Address inputs of t RCWS 10358's registers are thereby connected in parallel with address inputs of SR 10362's registers.

4 RAG 20288's address inputs to GRF 10354, and to BIAS 20246 and RCWS 10358, may select registers therein to be either source registers, that is registers providing I data, or destination registers, that is registers receiving data. RAG 20288's address outputs are designated as output Source and Destination Register Address (SDADR) of RAG 20288. RAG 20288's SDADR output is connected to address input of register comprising GRF 10354, BIAS 20246, and RCWS 10358. As described above, SRs 10362 are structured as 16 frames of 8 registers per ml--u---L~Y -388frame and RCWS 10358 is structured as a corresponding 16 frames of one register per frame. GRF 10354 and BIAS 20246 are structured and utilized as single registers but, for addressing purposes, are regarded as being comprised of 16 frames of 8 registers per frame. Each SDADR output of RAG 20288 is an 8 bit word wherein the most significant bit indicates whether the addressed register, either a Source or a Destination Register, reside in GRs 10360 or within SRs 10362, BIAS' 20246, and i10 RCWS 10358. The four next most significant bits comprise a frame select field for selecting one of 16 frames within GRs 10360 or within SRs 10362, BIAS 20246, and S. RCWS 10358. The three least significant bits comprise a register select field selecting a particular register within the frame selected by frame select field.

I Within a single system clock cycle, SDADR output of j RAG 20288 may select a source register and data may be read from that source register, or SDADR output may I 1 1select a destination register and data may be written into that destination register. As previously described, each JP 10114 microinstruction requires a minimum of twosystem clock cycles for execution, that is at first clock .I cycle in State MO and a second clock cycle in State Ml.

During a single microinstruction therefore, a source register may be selected and data read from thai source register, and a destination register selected and data written into that destination register. Certain operations, however, may require more than one -389microinstruction for execution. For example, a readmodify-write operation wherein data is read from a particular register, modified, and written back into that register may require two or more microinstructions for execution.

Referring first to RAG 20288 structure, RAG 20288 includes MISPR 10356. MISPR 10356 includes Top Of Stack Counter (TOSCNT) 25310, Top Of Stack-1 Counter (TOS-1CNT) 25312, and Bottom Of Stack Counter (BOSCNT) 25314.

10 Contents of TOSCNT 25310, TOS-ICNT, 25312 and BOSCNT 25314 are respectively, pointers to Current, Previous, and Bottom frames of SRs 10362, that is, to MIS 10368. As will be described below, these pointers are also utilized to address MOS 10370, TOSCNT 25310, TOS-1CNT 25312, and BOSCNT 25314 are each four bit binary counters comprised, for example, of SN74S163s.

Data inputs of TOSCNT 25310 to BOSCNT 25314 are connected from JPD Bus 10142. Control inputs of TOSCNT S% 15310 to BOSCNT 25314 are connected from microinstruction control outputs of FUSITT 11012. Data outputs of TOSCNT 25310 to BSOCNT 25314 are connected to data inputs of S Source Register Address Multiplexer (SRCADR) 25316 and to 1 04i data inputs of Destination Register Address Multiplexer (DSTADR) 25318. Data outputs of TOSCNT 25310 and BOSCNT 25314 are connected to inputs of Stack Event Monitor Logic (SEM) 25320.

Source and destination frame addresses are selected, as will be described further below, by SRCADR 25316 and 1 -390- DSTADR 25318 respectively. In addition to data inputs from TOSCNT 25310 and BOSCNT 25314, data inputs of SRCADR 25316 and DSTADR 25318 are connected from microinstruction word CONEXT subfield output from FUSITT 11012. Control inputs of SRCADR 25316 and DSTADR 25318 are connected from, respectively, microinstruction word RS and RD subfield outputs from FUSITT 11012. Source Frame Address Field (SRCFADR) output of.SRCADR 25316 and Destination Frame Address Field (DSTFADR) output of rr 10 DSTADR 25318 are connected to inputs of Source and Destination Register Address Multiplexer (SDADRMUX) 25322. SRCFADR and DSTFADR comprise frame select fields of RAG 20288, SDADR outputs for, respectively, source and destination registers.

In addition to SRCFADR and DSTFADR outputs of ADRSRC 25316 and DSTADR 25318, SDADRMUX 25322 receives microinstruction word SRC and DST subfield inputs from microinstruction outputs of FUSITT 11012. As previously described, SRC subfield is a 3 bit number designating a source register, that is, a source register within a frame selected by SRCFADR. DST is similarly a 3 bit number selecting a destination register within a frame indicated by DSTFADR. SRC subfield input to SDADRMUX 25322 is concatenated with SRCADR 25316 to respectively comprise, as described above, register and frame fields of a source register SDADR output of SDADRMUX 25322.

Similarly, DST subfield is concatenated with DSTFDADR output of DSTADR 25318 to comprise, respectively, register and frame subfields of a destination register SDADR output of SDADRMUX 25322. Selection between source 0 -391and destination register address inputs to SDADRMUX 25322, to generate a corresponding source of destination register SDADR output of SDADRMUX 25322 is controlled by microinstruction control inputs (not shown for clarity of presentation) connected to control inputs of SDADRMUX 25322. RDWS 25324 is a PROM decoding MD field from microinstruction words during reads from MEM 10112 and provides register select field of destination register rf address and selects one of the pointers as frame select r 10 field.

An Event output of SEM 25320 is connected to an input I' of EVENT 20284, previously described. SRCADR 25316, DSTADR 25318, and SDADRMUX 25322, as will be described further below, operate as multiplexers and may be comprised, for example, of SN74S153s.

Having described structure and organization of GRF 10354, BIAS 20246, and RCWS 10358, and structure of RAG 20288, operation of RAG 20288 to generate Source of L .Destination Register Address outputs SDADR will be described next below. Addressing of JP 10114's stack mechanism, comprising SRs 10362 and RCWS 10358, will be described first, followed by addressing of GRs 10360 and Ito# BIAS 20246.

SR 10362 portion of GRF 10354, RCWS 10358, and BIAS 20246 are addressed by Current, Previous, and Bottom Frame Pointers contained, respectively, in TOSCNT 25310, TOS-1CNT 25312, and BOSCNT 25314. Current, Previous, and Bottom Pointers comprise frame select fields of SDADRMUX 25322. As previously described, Current, Previous and Bottom Pointer outputs of TOSCNT 25310 to BOSCNT 25314 11\ 1 -392are provided as inputs of SRCADR 25316 and DSTADR 25318.

Microinstruction word RS subfield to control input of SRCADR 25316 selects either Current, Previous or Bottom Pointer input of SRCADR 25316 to comprise SRCFADR output of SRCADR 25316, that is to be frame select field of source register address. Similarly, microinstruction word RD subfield to control input of DSTADR 25318 V concurrently selects either Current, Previous, or Bottom Pointer inp'ts of DSTADR 25318 to comprise DSTADR 25318's concurrently selects either Current, Previous, or Bottom Pointer inputs of DSTADR 25318 to comprise DSTADR 25318s 1 DSTFADR output, that is frame select field of destination register address. As described above, SRCFADR and DSTFADR are provided as inputs to SDADRMUX 25322.

Microinstruction word SRC and DST subfield inputs to SDADRMUX 25322 concurrently determine, respectively, source and destination registers within source and destination frames specified by SRCFADR and DSTFADR.

SDADRMUX 25322 then, operating under microinstruction control, selects either SRCFADR and SRC to comprise SDADR output to SR 10362 as a source register address or selects DSTFADR and DST as SDADR output specifying a destination register address. By microinstruction control of SRCADR 25316, DSTADR 25318, and SDADRMUX 25322, a CS 10110 microprogram may select a source frame and register within SR 10362 and simultaneously specify a possible different destination frame and register within SR 10362, All possible combinations of source framit. and -393register and destination frame and register in GRs 10360, SRs 10362, BIAS 20246, and RCWS 10358 are valid.

Control of SRCADR 25316, DSTADR 25318, and SDADRMUX 25322 in addressing SR 10362 portion of GRF 10354, and RCWS 10358, is controlled, in part, by current CS 10110 state. Pertinent CS 10110 operating states, previously described, are State Ml and State RW. When CS 10110 is :in neither State RW nor State M1, SR 10362 is addressed through SRCADR 25316 and microinstruction word SRC subfield, that is SR 10362 and RCWS 10358 are provided t with source register addresses when CS 10110 is in "*neither RW nor M1 States. When CS 10110 enters State Ml, SR 10362 and RCWS 10358 is addressed through DSTADR 25318 and by microinstruction word DST subfield. That is, SR 15 10362 and RCWS 10358 are provided with destination register addresses during State Mi. Similarly, SR 10362 and RCWS 10358 are provided with destination register H addre5ses when CS 10110 is operating in State RW, that is when data is being read from MEM 10112 and written into SR 10362 or RCWS 10358. In this case, however, low order 3 bits of destination register address, that is register select field, are provided by RDS 25324, which decodes microinstruction word subfield MD (Memory Destination).

RDS 25324 also provides a control input that DSTADR 25318 to select one of Current, Previous, or Bottom Pointers C from MISPR 10356 to comprise frame select field of destination register address.

4 -394- As stated above, frame select field of source and destination register addresses are provided from TOSCNT 25310, TOS-1CNT 25312, and BOSCNT 25314. As described above, the most'significant bit of source and destination register address are forced to logic 1 or logic 0, depending upon whether GR 10360 or SR 10362, BIAS 20246, and RCWS 10358 are being addressed. Contents of TOSCNT 25310 to BOSCNT 25314, that is Current, Previous, and Bottom Pointers, are controlled by microinstruction Se 10 control outputs of FUSITT 11012. Current and Previous n Pointers change as stacks are "pushed" or "popped" to and from MIS 10368 as JP 10114 performs, respectively, calls ,t t and returns. Similarly, Current, Previous and Bottom Pointers will be incremented or decremented as MIS 10368 frames are transferred between MIS 10368 and MEM 10112, as previously described with respect to CS 10110's Stack Mechanisms.

Referring first to Current and Previous Pointer operation, Current and Previous Pointers in TOSCNT 25310 and TOS-1CNT 25312 are initially set, respectively, to point to Frames 1 and 0 of MIS 10368 by being loaded from SJPD Bus 10142. TOSCNT 25310 and TOS-1CNT 25312 are enabled to count when two conditions are met. First condition is dependent upon current operating state of CS 10110. TOSCNT 25310 and TOS-1CNT 25312 will be enabled to count during last system clock cycle of CS 10110 operating States Ml or AB. Second condition is dependant upon whether JP 10114 is to execute a call or return.

TOSCNT 25310 and TOS-1CNT 25312 may be enabled to count if a current microinstruction indicates JP 10114 is to r't -395execute a call or return, or if CS 10110 is exiting State AB as exit from State AB is an implied call operation.

Both a call and an implied call, that is exit from State AB, will cause TOSCNT 25310 and TOS-lCNT 25312 to be incremented. A return will cause TOSCNT 25310 and TOS- 1CNT 25312 to be decremented.

Referring to BOSCNT 25314, Bottom Frame Pointer is initially loaded from JPD Bus 10142 to point to MIS 10368 'Frame 1. Again, incrementing or decrementing of BOSCNT 25314 is dependant upon CS 10110 operating state and Soperation to be performed. BOSCNT 25314 is enabled to count upon exiting from State Ml. In addition, DEVCMD subfield of a current microinstruction word must indicate that BOSCNT 25314 is to be incremented or decremented.

BOSCNT 25314 will be incremented or decremented upon exit from State M1 as indicated by microinstruction word S" DEVCMD subfield.

SEM 25320 monitors relative values of Current and Bottom Pointers residing in TOSCNT 25310 and BOSCNT 25314 and provides outputs to EVENT 20284 for purposes of controlling operation of MI 10368 and MOS 10370. SEM 25320 is comprised of a Read Only Memory, for example '.93S427s, receiving Current and Bottom Pointers as inputs. SEM 25320 detects 3 Events occurring in operation of TOSCNT 25310 and BOSCNT 25314, and thus in operation of MIS 10368 and MOS 10370. First, SEM 25320 detects an MIS 10368 Stack Overflow. This Event is indicated if the present value of Current Frame Pointer is greater than 8 larger than the present value of Bottom I4 -396- Frame Pointer. Second, SEM 25320 detects when MIS 10368 contains only one frame of information. This event is indicated if the value of Current Frame Pointer is equal to the value of Bottom Frame Pointer. In this case, the previous frame of MIS 10368 resides in MEM 10112 and must be fetched from MEM 10112 before a reference to the previous stack frame may be made. Third, SEM 25320 detects when MIS 10368 and MOS 10370 are full. This Event is indicated if the present value of Current Frame I 10 Pointer is 16 larger than the present value of Bottom S Frame Pointer When this Event occurs, any further S b:rt attempt to write a frame onto MIS 10368 or MOS 10370 will I result in a MOS 10370 Stack Overflow. EVENT 20284 responds to these Events indicated by SEM 25320 by initiating execution of an appropriate Event Handling microinstruction sequence, as previously decribed. It should be noted that MIS 10368 and MOS 10370 are addressed in the same manner, that is through use of l Current, Previous and Bottom Frame Pointers and certain microinstruction word subfields. Primary difference *ibetween operation of MIS 10368 and'MOS 10370 is in the manner in which stack overflows are handled. In the case of MIS 10368, stack frames are transferred between MIS 10368 and MEM 10112 so that MIS 10368 is effectively a bottomless stack. MOS 10370, however, contains a maximum of 8 stack frames, in a present embodiment of CS 10110, so that no more than eight Events may be pushed onto MOS 10370 at a given time.

-397- GR 10360 is addressed in a manner similar to SR 10362, BIAS 20246, and RCWS 10358, that is through ADRSRC 25316, DSTADR 25318, and SDADRMUX 25322. Again, register select fields of source and destination register addresses are provided by microinstruction word SRC and DST subfields. Frame select field of source and destination register addresses is, however, specified by microinstruction word CONEXT subfield. -In this case, S. microinstruction word RS and RD subfields specify that S 10 frame select fields of source and destination register addresses are to be provided by CONEXT subfield.

Accordingly, ADRSRC 25316 and DSTADR 25318 provide CONEXT subfield as SRCFADR and DSTFADR inputs to SDADRMUX 25322.

Having described structure and operation of RAG 20288, TIMERS 20296 will be described next below.

Referring to Fig. 254, a partial block diagram of TIMERS 20296 is shown. As indicated therein, TIMERS 20296 includes Interval Timer (INTTMR) 25410, Egg Timer (EGGTMR) 25412, and Egg Timer Clock Enable Gate (EGENB) 25416.

d.d.d. Times--20-296--+f4i--254 S" Referring first to INTTMR 25410, a primary function of INTTMR 25410 is to maintain CS 10110 architectural time as previously described with reference to Fig. 106A and previous descriptions of CS 10110 UID addressing. As described therein, a portion of all UID addresses generated by all CS 10110 systems is an Object Serial Number (OSN) field. OSN field uniquely defines each -398object created by operation of or for use in a particular CS 10110. OSN field of an object's UID is, in a particular CS 10110, generated by determining time of creation of that object relative to an arbitrary historic starting time common to all CS 10110 systems. That time is maintained within a MEM 10112 storage space, or address location, but is measured by operation of INTTMR 25410.

e INTTMR 25410 is a 28 bit counter clocked by a 110 Nano-Second Clock (11ONSCLK) input and is enabled to i* count by a one MHZ Clock Enable input (CLK1MHZENB).

INTTMR 25410 may thereby be clocked at a one MHZ rate to measure one microsecond intervals. Maximum time interval which may be measured by INTTMR 25410 is thereby 268.435 seconds.

As indicated in Fig. 254, INTTMR 25410 may be loaded from and read to JPD Bus 10142. In normal operation, the t MEM 10112 location containing architectural time for a particular CS 10110 will be loaded with current architectural time at time of start up of that particular CS 10110. INTTMR 25410 will concurrently be loaded with all zeros. Thereafter, INTTMR 25410 will be clocked at one microsecond intervals. Periodically, when INTTMR 25410 overflows, architectural time stored in MEM 10112 will be accordingly updated. At any time, therefore, current architectural, time may be determined, down to a one microsecond increment, by reading architectural time from the previous updated architectural time stored in -399- MEM 10112 and elapsed interval since last update of architectural time from INTTMR 25410. In the event of a failure of CS 10110, architectural time in MEM 10112 and INTTMR 25410 may be saved in MEM 10112 by reading elapsed intervals since last architectural time update. When normal CS 10110 operation resumes, INTTMR 25410 may be reloaded with a count reflecting current architectural time. As indicated in Fig. 254, INTTMR,25410 is loaded from JPD Bus 10142 when INTTMR 25410 is enabled by a Load 10 Enable input (LDE) provided from DP 10118.

Referring to EGGTMR 25412, certain CS 10110 Events, in particular Asychronous Events previously described S' with reference to EVENT 20284, are received or acknowledged by EVENT 20284 only at conclusion of State Ml of first microinstruction of an SOP. As certain CS 10110 microinstructions have long execution times, these Asynchronous Events may be subjected to an extended latency, or waiting, interval before being serviced.

EGGTMR 25412, in effect, measures latency time of pending Asychronous Events and provides an output to EVENT 20284 if a predetermined maximum latency time is exceeded.

As indicated in Fig. 254, EGGTMR 25412 is clocked by a 110 Nano-Second Clock input (11ONSCLK) EGGTMR 25412 is initially set to zero by load input (LDZRO) at end of State Ml of the first microinstruction of each SOP executed by CS 10110, or when specifically instructed so by DEVCMD subfield of a micrc nstruction word. EGGTMR 25412 is incremented when enabled by Clock Enable (CLKENB) input from EGGENB 25416. There are two conditions necessary for EGGTMR 25412 to be incremented.

-400- First condition is occurrence of an Asychronous Event, which is indicated by input ASYEVNT to EGGENB 25416 from EVENT 20284 Second condition is that 16 or more microseconds have elapsed since last increment of EGGTMR 25412. This interval is measured by an output from fourth bit of INTTMR 25410 which, as shown in Fig. 254, is connected to an input of EGGENB 25416. EGGTMR 25412 iis a four bit counter and will thereby overflow and generate output OVRFLW to EVENT 20284 256 microseconds 10 after beginning of an SOP if an Asychronous Event has S, occurred and if at least 16 microseconds have elapsed since start of that SOP EGGTMR 25412 thereby insures a maximum service latency of 256 microseconds for Asychronous Events.

e.e.e. Fetch--Unt--10-2 2 -T nterface--to--Becute Uniti0l2-2 Finally, as previously described FU 10120's interface to EU 10122 is primarily comprised of EUDIS Bus 20206, for providing EUDPs to EU 10122's EUSITT, and FUINT 20298. Operation of EUSDT 20266 and EUDIS Bus 20206 has t t b I been previously described and will be described further in a following description of EU 10122. FUINT 20298 is primarily concerned with generating Event Requests for conditions signalled from EU 10122 so that these Events may be serviced. In this regard, FUINT 20298 is primarily comprised of gates receiving Event Requests from EU 10122 and providing corresponding outputs to EVENT 20284. Another interface function performed by FUINT 20298 is generation of a "transfer complete" signal -401generated by FU 10122 and provided to EU 10122 to assert that a EU 10122 result read from EU 10122 to FU 10120 has been received. This transfer complete signal indicates to EU 10122 that EU 10122's result register, described in a following description.of EU 10122, is available for further use by EU 10122. This transfer complete signal is generated by an output of FUSITT 11012 as part of microinstruction sequences for transferring data from EU 10122 to FU 10120 or MEM 10112.

10 Having described structure and operation of FU 10120, including DESP 20210, MEMINT 20212, and FUCTL 20214, the t f: structure and operation of EU 10122 will be described C" next below.

C. E uxecute- Untt--O-20-2--+P i-r-2-r-25-25 As previously described, EU 10122 is an arithmetic processor capable of executing integer, packed and *unpacked decimal, and single and double precision floating point arithmetic operations. A primary function of EU 10122 is to relieve FU 10120 of certain arithmetic operations, thus enhancing efficiency of CS 10110.

.Transfer of operands from MEM 10112 to EU 10122 is controlled by FU 10120, as is transfer of results of i*4 arithmetic operations from EU 10122 to FU 10120 or MEM 10112. In addition, EU 10122 operations are initiated by FU 10120 by EU 10122 Dispatch Pointers invited to EU 10122 by EUSDT 20266. EU 10122 Dispatch Pointers may initiate both arithmetic operations required for i i ~1 r 1 -402execution of SINs and certain EU 10122 operations assisting in handling of CS 10110 events. As previously described, EU 10122 Dispatch Pointers are translated into sequences of microinstructions for controlling EU 10122 by EU 10122's EUSITT which is similar in structure and operation to FUSITT 11012. As will be described further below, EU 10122 includes a command queue for receiving A and storing sequences of EU 10122 Dispatch Pointers from FU 10120. In addition, EU 10122 includes a general S 10 register file, or scratch pad memory, similar to GRF 1 10354. EU 10122's general register file is utilized, in part, in EU 10122 Stack Mechanisms similar to FU 10120's SR's 10362.

Referring to Fig. 203, a partial block diagram of EU 10122 is shown. EU 10122's general structure and 'operation will be described first with referene to Fig.

203. Then EU 10122's structure and operation will be described in further detail with aid of subsequent figures which will be presented as required.

As indicated in Fig. 203, major elements of EU 10122 include Execute Unit Control Logic (EUCL) 20310, Execute Unit I1 Buffer (EUIO) 20312, Multiplier Logic (MULT) 20314, Exponent Logic (EXP) 20316, Multiplier Control Logic (MULTCNTL) 20318, and Test and Interface Logic (TSTINT) 20320. EUCL 20310 receives Execute Unit Dispatch Pointers (EUDP'S) from EUSDT 20266 and provides corresponding sequences of microinstructions to control operation of EU 10122.

-403- EUIO 20312 receives operands, or data, from MEM 10112, translates those operands into certain formats most efficiently used by EU 10122. EUIO 20312 receives results of EU 10122's operations and translates those results into formats to be returned to MEM 10112 or FU 10120, and presents those results to MEM 10112 and FU 10120.

MULT 20314 and EXP 20316 are arithmetic units for ,performing arithmetic manipulations of EU 10122 S' 10 operations. In particular, EXP 20316 performs operations with respect to exponent fields of single and double precision floating point operations. MULT 20314 performs arithmetic manipulations with respect to mantissa fields of single and double precision floating point operations, and arithmetic operations with regard to integer and packed decimal operations. MULTCNTL 20318 controls and coordinates operation of MULT 20314 and EXP 20316 and Sprealignment and normalization of mantissa and exponent I fields in floating point operations. Finally, TSTINT 20320 performs certain test operations with regard to EU a, t 10122's operations, and is the interface between EU 10122 and FU 10120.

a. senercutent-of/- £-2 1 f ectlte^Un+^-OQ-M- 2 2 Referring first to EUIO 20312, EUIO 20312 includes Operand Buffer (OPB) 20322, Final Result Output Multiplexer (FROM) 20324, and Exponent Output Multiplexer (EXOM) 20326. OPB 20322 has first and second inputs connected, respectively, from MOD Bus 10144 and JPD Bus 10142. OPB 20322 has a first output connected to a first Pinr zi~s M 1, i r_ 11111~ -404input of Multiplier Input Multiplexer (MULTIM) 20328 and MULT 20314. A second output of OPB 20322 is connected to first inputs of Inputs Selector A (INSELA) 20330 and Expoitent Execute Unit General Register File Input Multiplexer (EXRM) 20332 in EXP 20316.

FROM 20324 has an output connected to JPD Bus 10142.

A first input of FROM 20324 is connected from output of at Multiplier Execute in General Register File Input Multiplexer (MULTRM) 20334 and MULT 20314. A second input of FROM 20324 is connected from output of Final Result Register (RFR) 20336 of MULT 20314. EXOM 20326 has an output conne-ted to JPD Bus 10142. EXOM 20326 is a first input connected from output of Scale Register (SCALER) 20338 of EXP 20316. EXOM 20326 has second and third inputs connected from outputs of, respectively, Next Address Generator (NAG) 20340 and Command Queue (COMQ) 20342 of EUCL 20310.

S. 2. Execute-Unit--eontro Loai-2-31-0 i Referring to EUCL 20310, EUCL 20310 includes NAG 20340, COMQ 20342, Estate Unit S Interpreter Table (EUSITT) 20344, and Mitccinstruction Control Register and Sf Decode Logic (mCRD) 20346. COMQ 20342 has an input connected from EUDIS Bus 20206 for receiving SDPs from EUSDT 20266. COMQ 20342 has, as described above, a first output connected to a third input of EXOM 20326, and has a second output connected to an input of NAG 20340. NAG 20340 has, as described above, a first output connected to second input of EXOM 20326. NAG 20340 has a second output connected to a first input of EUSITT 20344. As previously described, EUSITT 20344 corresponds to FUSITT -405- 91 4 #9 9 9, t 3 3 C1 I. 3 Itr L 41 CI II 99**e 34 9.

~0 11012 and stores sequences of microinst:uctions for controlling operation of EU 10122 in response to EU 10122 Dispatch Pointers from FU 10120. EUSITT 20344 has a second input connected from JPD Bus 10142 and has an output connected to input of mCRD 20346. mCRD 20346 includes a register and logic for receiving and decoding microinstructions provided by EUSITT 20344. In addition to an input from EUSITT 20344, mCRD 20346 has first outputs providing decoded microinstruction control signals to all parts of EU 10122. mCRD 20346 also has a second output connected to a first input of Input Selecter B (INSELB) 20348 and EXP 20316.

3. Mutip'lexer-L-o-cc-2f-4 Referring to MULT 20314, MULT 20314 includes two 15 parallel arithmetic operation paths for performing addition, subtraction, multiplication, and division operations on packed decimal numbers, integer numbers, and mantissa portions of single and double precision floating point numbers. MULT 20314 also includes a related portion of EU 10122's general register file, a memory for storing constants used in arithmetic operations, and certain input data selection circuits.

That portion of EU 10122's GRF residing in MULT 20314 is comprised of Multiplier Register File (MULTRF) 20?.0.

Output of MULTRF 20350 is connected to a second input of MULTIM 20328. A first input of MULTRF 20350 is connected from output of RFR 20336 and a second input of MULTRF 20350 is connected from output of MULTRM 20334. First j i.

-406and second inputs of MULTRM 20334 are in turn connected, respectively, from output of RFR 20336 and from output of Container Size Logic (CONSIZE) 20352 of TSTINT 20320.

MULTIM 20328 selects the data inputs to MULT 20314's arithmetic circuits and has, as previously described, first and second inputs connected respectively from first output of OPB 20322 and from output of MULTRF 20350.

Output of MULTIM 20328 is connected through Multiplier (MULT) Bus 20354 to input of Multiplier Quotient Register (MQR) 20356 and to input of Nibble Shifter (NIBSHF) 20358. Another input to MQR 20356 and NIBSHF 20358 is provided by Constant Store (CONST) 20360. CONST 20360 is a memory for storing constant values used n MULT 20314 operations. Output of CONST 20360 is connected to MULT Bus 20354. MULT 20314's arithmetic circuits may thereby be provided with inputs from OPB 20322, MULTRF 20350, and CONST 20360.

SOMULT 20314's arithmetic circuitry is comprised of two, parallel arithmetic operation paths having, as 20 common inputs, outputs of MULTIM 20328 and CONST 20360.

Common termination of these parallel arithmetic operation paths is Final Register Shifter (FRS) 20362. A first arithmetic operation path is provided through NIBSHF 20358, whose input is connected from MULT Bus 20354.

NIBSHF 20358's output is connected to a first input of FRS 20362 and a control input of NRBSHF 20358 is connected from an output of Multiplier Control Logic (MULTCNT) 20364 and MULTCNTL 20318,, 1 1 -407- MULT 20314's second arithmetic operation path is provided through MQR 20356. As described above, MQR 20356's input is connected from MULT Bus 20354. MQR 29356's output is connected to first and second inputs of Times 1 And Times 2 Multiply Shifter (MULTSHFT12) 20366 and Times 4 And Times 8 Multiply Shifter (MULTSHFT48) 20368. Outputs of MULTSHFT12 and MULTSHFT8 are o connected, respectively, to first and second inputs of First Multiplier Arithmetic and Logic Unit (MULTALUI) S 10 20370. MULTALU1 20370's output is connected to input of Multiplier Working Register (MWR) 20372. Output of MWR 20372 is connected to a first input of Second Multiplier Arithmetic and Logic Unit (MULTALU2) 20374. A second input of MULTALU2 20374 is connected from output of RFR 20336. Output of MULTALU2 is connected to a second input of FRS 20362. As described above, first input of FRS 20362 is connected from output of NIBSHF 20368. Output t of FRS 20362 is connected to input of RFR 20336.

As described above, output of RFR 20336 is connected to second input of MULTALU2 20374, to first input of MULTRF 20350, to first input of MULTRM 20334, and to second input of FROM 20324. Output of RFR 20336 is also connected to input of Leading Zero Detector (LZD) 20376 of MULTCNTL 20318, and to inputs of Exception Logic (ECPT) 20378, CONSIZE 20352, and TSTINT 20320.

4. Exponent-Lo ic-2-2326 Referring to EXP 20316, as previously described EXP 20316 performs certain operations with respect to exponent fields of single and double precision floating point number in EU 10122 floating point operations. EXP 3A -408-

C

t Sb 4 5 i r Ir S C. I t C~ I I 6 C Crt C O I

CICE

SI,,

I b* Cs C 4 20316 includes a second portion of EU 10122's general register file, shown herein as Exponent Register File (EXPRF) 20380. Although indicated as individual register files, MULTRF 20350 and EXPRF 20380 comprise, as in GRF 10354, a unitary register file structure with common, parallel addressing of corresponding registers therein.

Output of EXPRF 20380 is connected to a second input of INSELA 20330. A first input of EXPRF 20380 is connected from output of EXRM 20332. As previously described, a first input of EXRM 20332 is connected from second output of OPB 20322 through EXPQ Bus 20325. A second input of EXRM 20332 is connected from output Scale Register (SCALER) 20338. A second input of EXPRF 20380 is connected from output of Sign Logic (SIGN) 20382.

Input of SIGN 20382 is connected from second output of SCALER 20338.

INSELA 20330, INSELB 20348, Exponent ALU (EXPALU) 20384 and SCALER 20338 comprise EXP 20316's arithmetic circuitry for manipulating exponent fields of floating 20 point numbers. INSELA 20330 and INSELB 20348 select, respectively, first and second inputs to EXPALU 20384.

As previously described, a first input of INSELA 20330 is connected from second output of OPB 20322 through EXPQ Bus 20325. Second input of INSELA 20330 is connected from output of EXPRF 20380. Output of INSELA 20330 is connected to first input of EXPALU 20384. First input of INSELB 20348 is, as previously described, connected from a second output of mCRD 20346. Second input of INSELB *4

'I

4t -409- 20348 is connected from output of OPB 20322 through EXPQ Bus 20325. Third input of INSELB 20348 is connected from output of SCALER 20338 and fourth input of INSELB 20348 C is connected from output of LZD 20376. Output of INSELB 20348 is connected to second input of EXPALU 20348.

Output of EXPALU 20348 is connected to input of SCALER 20338.

As previously described, second output of SCALER 20338 is connected with input of SIGN 20382 and first 10 output is connected to second input of EXRM 20332 and to third input of INSELB 20348. First output of SCALER 20338 is also connected to EXPQ Bus 20325, to first input of EXOM 20326, and to a second input of MULTCNT 20364.

Mu-lt p-ie---Crontror--2-31-8 As previously described, MULTCNTL 20318 provides certain control signals and information for controlling and coordinating operation of EXP 20316 and MULT 20314 in performing arithmetic operations on floating point numbers. MULTCNTL 20318 includes LZD 20376 and MULTCNT 20364. Input of LZD 20376 is connected from output of RFR 20336 through FR Bus 20337. Output of LZD 20376 are connected,to a second input of MULTCNT 20364 and to fourth input of INSELB 20348. A second input of MULTCNT 20364 is connected from output of SCALER 20338. As previously described, control output of MULTCNT 20364 is connected to control inputs of NIBSHF 20358.

6. ests and-.nteTace--oo4ce-26-10 Finally, TSTINT 20320 includes ECPT 20378, CONSIZE 20352, and Testing Condition Logic (TSTCON) 20386. Input of ECPT 20378 and first input of CONSIZE 20352 are -410connected from output of RFR 20336 through FR Bus 20337.

A second input of CONSIZE 20352 is connected from LENGTH Bus 20226. An output of CONSIZE 20352 is connected, together with other inputs from EU 10122 (not shown for clarity of presentation) to TSTCON 20386. Output of TSTCON 20386 (not shown for clarity of presentation) are connected to NAG 20340. TSTCON 20386 and ECPT 20378 have outputs to and inputs from FU 10120's FUINT 20298.

s Having described the overall structure of EU 10122 10 above, operation of EU 10122 will be described next below with aid of further diagrams which will be introduced as required. Finally, operation of TSTINT 20320 will be described, including a description of the detailed control signal interface between EU 10122 and FU 10120 S 15 through TSTINT 20320 and FUINT 20298. In addition to Sadefining the interface between EU 10122 and FU 10120, certain features of EU 10122 operation will be described C V wherein those operations are executed in cooperation with MEM 10112 and FU 10120. For example, EU 10122's Stack Mechanisms, comprising in part portions of MULTRF 20350 Sand EXPRF 20380, resides partly in MEM 10112 so that L. operation of EU 10122's Stack Mechanisms requires cooperative operations by EU 10122, MEM 10112 and FU 10120.

b. xece-Unit-- G---ope- 1 5 xrecuge--Un TSTIN 20320 a F Referring to Fig. 255, a more detailed block diagram of EUCL 20310 is shown. As described above, EUCL 20310 «4.

w eren th se pertio s ar ex cutd i cooeraionwit MEM 011 andFU 0120 Fo exmple EU1012's tac -411receives EU 10122 Dispatch Pointers through EUDIS Bus 20206 from EUSDT 20266 and FUCTL 20214. EU 10122 Dispatch Pointers select certain EU 10122 microinstruction sequences for executing EU 10122 arithmetic operations as required to execute user's programs, that is SOPs, and to assist in handling JP 10114 Events. As described above, major elements of EUCL *20310 include COMQ 20342, EUSITT 20344,,mCRD 20346, and NAG 20340.

e 10 a.a. Command--Otette--2034-2 I Inputs of COMQ 20342 are connected from EUDIS Bus S. 20206 to receive and store EU 10122 Dispatch Pointers provided from EUSDT 20266. Each such EU 10122 Dispatch Pointer is comprised of two information fields. A first information field contains a 10 bit starting address of a corresponding sequence of microistructions residing in EUSITT 20344. Second field of each EU 10122 Dispatch Pointer is a 6 bit field containing certain control information, such as information identifying data format of corresponding operands to be operated upon. In this 'case unit dispatch pointer control field bits specify whether operands to be operated upon comp-ise signed or unsigned integer, packed or unpacked decimal, or single or double precision floating point numbers.

COMQ 20342 is comprised of two one word wide by two word deep register files. A first of these register fields is comprised of SOP Command Queue Control Store (CQCS) 25510 and SOP Command Queue Address Store (CQAS) f -412- 25512. Together, CQCS 25510 and CQAS 25512 comprise a one word wide by two word deep register file for receiving and storing EU 10122 Dispatch Pointers corresponding to SOPs, that is Dispatch Pointers for initiating EU 10122 operations directly concerned with executing a user's program. Address fields of these SOPs are received in CQAS 25512, while control fields are received and stored in CQCS 25510. COMQ 20342 is thereby capable of receiving and storing up to two sequential EU 10 10122 Dispatch Pointers corresponding to user program i .SOPs These SOP derived Dispatch Pointers are executed in the order received from FU 10120. EU 10122 is thereby capable of receiving and storing one currently executing SOP Dispatch Pointer and one pending SOP Dispatch Pointer. Further SOP Dispatch Pointers may be read into COMQ 20342 as previous SOPs are executed.

Sb.b. eommand-Oueue--ven-ez~re-Ste -255 4--and--eommand-8ef -en CmAddee en t otro tro-- e (Q 2551 Command Queue Event Control Store (CQCE) 25514 and command Queue Event Address Control Store (CQAE) 25516 are similar in function and operation to, respectively, CQCS 25510 ad CQAS 25512. CQCE 25514 and CQAE 25516 I receive and store, however, EU 10122 Dispatch Pointers initiating EU 10122 operations requested by FU 10120 as required to handle JP 10114 Events. Again, CQCE 25514 and CQAE 25516 comprise a one word wide by two word deep register file. CQAE 25516 receives and stores address fields of Event Dispatch Pointers, while CQCE 25514 receives and stores corresponding control fields of Event -413- Dispatch Pointers. Again, COMQ 20342 is capable of receiving and storing up to two sequential Event Dispatch Pointers at a time.

As indicated in Fig. 255, outputs of CQAS 25512 and CQAE 25516, that is address fields of EU 10122 Dispatch Pointers are provided as inputs to Select Case Multiplexer (SCASE) 25518 and Starting Address Select Multiplexer (SAS) 25520 and NAG 20340, which will be Sdescribed further below. Control field outputs of CQCS 10 25510 and CQCE 25514 are provided as inputs to OPB 20322, described further below.

c.c. Execute--Untt-S-Interpreter--Tbie 20344 Referring to EUSITT 20344, as described above EUSITT 20344 is a memory for storing sequences of Et microinstructions for controlling operation of EU 10122 in response to EU 10122 Dispatch Pointers received from f A FU 10120. These microinstruction sequences may, in general, direct operation of EU 10122 to execute S 20 arithmetic operations in response to SOPs of user's S•programs, or aid direct execution of EU 10122 operations required to service JP 10114 Events. EUSITT 20344 may be, for example, a 60 bit wide by 1,280 word long memory structured as pages of 128 words per page. A portion of EUSITT 20344's pages may be contained in Read Only Memory, for example for storing sequence of microinstructions for handling JP 10114 Events.

Remaining portions of EUSITT 20344 may be ccnstructad of I' ii s -414- Random Access Memory, for example for storing sequences of microinstructions for executing EU 10122 operations in response to user program SOPs. This structure allows EU 10122 microinstruction sequences concerned with operation of JP 10114's internal mechanisms, for example handling of JP 10114 Events, to be effectively permanently stored 1 .in EUSITT 20344. That portion of EUSITT 20344 t constructed of Random Access Memory may.be used to store sequences of microinstructions for executing SOPs. These :v 10 Random Access Memories may be used as writable control store to allow sequences of microinstructions for executing SOPs of one or more S-Languages currently being utilized by CS 10110 to be written into EUSITT 20344 from MEM 10112 as required.

As previously described, EUSITT 20344's second input is a Data (DATA) input connected from JPD Bus 10142.

EUSITT 20344's data input is utilized to write sequences of microinstructions into EUSITT 20344 from MEM 10112 through JPD Bus 10142. EUSITT 20344's first input is an address (ADR) input connected from output of Address S* Driver (ADRD) 25522 and NAG 20340. Address inputs S" provided by ADRD 25522 select word locations within EUSITT 20344 for writing of microinstructions into EUSITT 20344, or for reading of microinstructions from EUSITT 20344 to mCRD 20346 to control operation of EU 10122.

Generation of these address inputs to EUSITT 20344 by NAG 20340 will be described further below.

d.d. Microcode-Geontrol-DecOde--Regte -2Q0-346 Output of EUSITT 20344 is connected to input of mCRD 20346. As previously described, mCRD 20346 is a register P. -2- -415- Vr

V

r, V t I 4*r 4 44 4 for receiving microinstructions from EUSITT 20344, and decoding logic for decoding those microinstructions and providing corresponding control signals to EU 10122. As indicated in Fig. 255, Diagnostic Processor Micro-Program Register (DPmR) 25524 is a 60 bit register connected in parallel with output of EUSITT 20344 to input of mCRD 20346. DPmR 25524 may be loaded with 60 bit microinstructions by DP 10118. Diagnostic microinstructions may thereby be provided directly to input of mCRD 20346 to provide direct microinstruction by microinstruction control of EU 10122.

Outputs of mCRD 20346 are provided, in general, to all portions of EU 10122 to control detailed operations of EU 10122. Certain outputs of mCRD 20346 are connected to inputs of Next Address Source Select Multiplexer (NASS) 25526 and Long Branch Page Address Gate (LBPAG) 25528 and NAG 20340. As will be described further below, these outputs of mCRD 23046 are used in generating address inputs to EUSITT 20344 when particular microinstructions sequences call for Jumps or Long Branches to other microinstruction sequences. Outputs of mCRD 20346 are also connected in parallel to inputs of Execution Unit Micro-Instruction Parity Check Logic (EUmIPC) 25530. EUmIPC 25530 checks parity of all microinstruction outputs of mCAD 20346 to detected errors in mCRD 20346's outputs.

e.e. Next -Address-Gaenerater--242-4 As described above, read and write addresses to EUSITT 20344 provided by NAG 20340 through ADRD 25522.

u

I

-416- Address inputs to ADRD 25522 are provided from either NASS 25526 or Diagnostic Processor Address Register (DPAR) 25532. In normal operation, address inputs to EUSITT 20344 are provided from NASS 25526 as will be described momentarily. DP 10118, however, may load EUSITT 20344 addresses into DPAR 25532. These addresses may then be read from DPAR 25532 through ADRD 25522 to individually select address locations within EUSITT 20344. DPAR 25532 may be utilized, in particular, to 10 provide addresses to allow stepping through of EU 10122 microinstruction sequences microinstruction by microinstruction.

As described above, NASS 25526 is a multiplexer having inputs from three NAG 20340 address sources. NASS 1 15 25526's first address input is from Jump (JMP) output of S, mCRD 20346 and LBPAG 25528. These address inputs are utilized, in part, when a current microinstruction calls for a Jump or Long Branch to another microinstruction or microinstruction sequence. Second address source is provided from SAS 25520 and, in general, is comprised of starting addresses of microinstruction sequences. SAS 25520 is a muliplexer having a first input from CQAS 25512 and CQAE 25516, that is starting addresses of microinstruction sequences corresponding to SOPs or for servicing JP 10114 Events. A second SAS 25520 input is provided from Sub-routine Return Address Stack (SUBRA) 25534. In general, and as will be described further below, SUBRA 25534 operates as a stack mechanism for storing current microinstruction addresses of interrupted microinstruction sequences. These stored addresses may -417subsequently be utilized to resume execution of those interrupted microinstruction sequences. Third address source to NASS 25526 is provided from Sequential and Case Address Generator (SCAG) 25536. In general, SCAG 25536 generates address to select sequential microinstructions within particular microinstruction sequences. SCAG 25536 also generates microinstruction address for microinstruction Case operations. As indicated in Fig.

255, outputs of SCAG 25536 and of SAS 25520 are bused 10 together to comprise a single NASS 25526 input.

Selection between outputs of SCAG 25536 and SAS 25520 are provided by control inputs (not shown for clarity of presentation) to SCAG 25536 and SAS 25520. Selection between NASS 25526's address inputs is controlled by Next Address Source Select Control Logic (NASSC) 25538, which provides control inputs to NASS 25526. NASSC 25538 is effectively a multiplexer receiving control inputs from TSTCON 20386 and TSTINT 20320. As will be described ffurther below, TSTCON 20386 monitors certain operating conditions or states within EU 10122 and provides corresponding inputs to NASSC 25538. NASSC 25538 effectively decodes these control inputs from TSTCON 20386 to provide selection control input to NASS 25526.

Having described overall structure and operation of NAG 20340, operation of NAG 20340 will be described in further detail next below.

Referring first to NASS 25526's address inputs provided from JMP output of mCRD 20346 and LBPAG 2!V- -418this address source is provided to allow selection of a next microinstruction by a current microinstruction. JMP output of mCRD 20346 allows a durrent microinstruction to direct a Jump to another microinstruction within the same page of EUSITT 20344. NASS 25526's input through LBPAG 25528 is provided from another portion of mCRD 20346's output specifying pages within EUSITT 20344. This input through LBPAG 25528 allows execution of Long Branch operations, that is jumps from a microinstruction in one page of EUSITT 20344 to a microinstruction in another St ,page. In addition, NASS 25526's input from JMP output of mCRD 20346 and through LBPAG 25528 is utilized to execute an Idle, or Standby, routine when EU 10122 is not currently executing a microinstruction sequence requested by FU 10120. In this case, Idle routine directs TSTCON 20386 to monitor EU 10122 Dispatch Pointer inputs to EU 10122 from FU 10120. If no EU 10122 Dispatch Pointers are present in COMQ 20342, or none are pending, TSTCON t' 20386 will direct NASSC 25538 to provide control inputs to NASS 25526 to select NASS 25526's input from mCRD 20346 and LBPAG 25528. Idle routine will continually test for EU 10122 Dispatch Pointer inputs until such a Dispatch Pointer is received into COMQ 20342. At this time, TSTCON 20386 will detect the pending Dispatch Pointer and direct NASS 25538 to provide control outputs to NASS 25526 to select NASS 25526's input from, in general, SAS 25520. TSTCOND 20386 and NASSC 25538 will also direct NASS 25526 to select inputs from SAS 25520 -419o e *4 4.

44.

C 4. 4 4-I $4 4..

4 *tr upon return from a called microinstruction to a previously interrupted microinstruction sequence,, As described above, SAS 25520 receives starting addresses from COMQ 20342 and from SUBRA 25534. SAS 25520 will select the output of CQAS 25512 or of CQAE 25516 as the input to NASS 25526 when a new microinstruction sequence is to be initiated to execute a user's program SOP or to service a JP 10114 Event. SAS 25520 will select an address output of SUBRA 25534 upon 10 return from a called sub-routine to a previously executing but interrupted sub-routine. SUBRA 25534, as described above, is effectively a stack mechanism for storing addresses of currently executing mcroinstructions when those microihstruction sequences are interrupted.

SUBRA 25534 is an 11 bit wide by 8 word deep register with certain registers dedicated for use in stacking Event Handling microinstruction sequences. Other portions of SUBRA 25534 are utilized for stacking of microinstruction sequences for executing SOPs, that is for stacking microinstruction sequences wherein a first microinstruction sequence calls for a second microinstruction sequence. SUBRA 25534 is not operated as a first-in-first out stack, but as a random access memory wherein address inputs selecting registers and SUBRA 25534 are provided by microinstruction control outputs of mCRD 20346. Operations of SUBRA 25534 as a stack mechanism is thereby controlled by the microinstruction sequences stored in EUSITT 20344. As indicated in Fig. 255, addresses of current 0 microinstructions of interrupted microinstruction i; j ~i 1 -420sequences are provided to data input of SUBRA 25534 from output of SCAG 25536, which will be described next below.

As described above, SCAG 25536 generates sequential addresses to select sequential microinstructions within microinstruction sequences and to generate microinstruction addresses for Case operations. SCAG 25536 includes Next Address Register (NXTR) 25540, Next Address Arithmetic and Logic Unit (NAALU) 25542, and SCASE 25518. NAALU 25542 is a 12 bit arithmetic and 10 logic unit. A first eleven bit input of NAALU 25542 is connected from output of ADRD 25522 and is thereby t current address provided to EUSITT 20344. A second four bit input to NAALU 25542 is provided from output of SCASE 25518. During sequential execution of a microinstruction sequence, output of SCASE 25518 is binary zeros and carry input of NAALU is forced to 1. Output of NAALU 25542 will thereby be and address one greater than the current microinstruction address provided to EUSITT 20344 and will thereby be the address of the next sequential microinstruction. As indicated in Fig. 255, SCASE 25518 receives an input from output of SCALER 20338. This ,input is utilized during Case operations and allows a data sensitive number to be selected as SCASE 25518's Soutput into second input of NAALU 25542. SCASE 25518's input from SCALER 20338 thereby allows NAG 20340 to perform microinstruction Case operations wherein Case Values are determined by the contents of SCALER 20338.

Next address outputs of NAALU 25542 are loaded into NXTR 25540, which is comprised of tri-state output registers. Next address outputs of NXTR 25540 are I*

I

-421connected, in common with outputs of SAS 25520, to second input of NASS 25526 as described above. During normal execution of microinstruction sequences, therefore, SCAG 25536 will, through NASS 25526 and ADRD 25522, select sequential microinstructions from EUSITT 20344. SCAG 25536 may also, as just described, provide next microinstruction addresses in microinstruction Case operations.

S 'f In summary, NAG 20340 is capable of performing all 10 usual microinstruction sequence addressing operations.

For example, NAG 20340 allows selection of next microinstructions by current microinstructions, either for Jump operations or Long Branch operations, through NASS 25525's input from mCRD 20346's JMP or through LBPAG 25528. NAG 20340 may provide microinstruction sequence starting addresses through COMQ 20342 and SAS 25520, or may provide return addresses to interrupted and stacked microinstruction sequences through SUBRA 25534 and SAS I .25520. NAG 20340 may sequentially address microinstructions of a particular microinstruction sequence through operation of SCAG 25536, or may perform :mciroinstruction Case operations through SCAG 25536.

H L 2. Operand--Btrfer--2M-2-2 Having described structure and operation of EUCL 20310, structure and operation of OPB 20322 will be described next below. As previously described, OPB 20322 receives operands, that is data, from MEM 10112 and FU 10120 through MOD Bus 10144 and JPD Bus 10142. OPB 20322 may then perform certain operand format translations to provide data to MULT 20314 and EXP 20316 in the formats

J,

-422most efficiently utilized by MULT 20314 and EXP 20316.

As previously described, EU 10122 may perform arithmetic operations on integer, packed and unpacked decimal, and single or double precision floating point numbers.

In summary, therefore, OPB 20322 is capable of accepting integer, single and double precision floating point, and packed and unpacked decimal operands from MEM 10112 and FU 10120 and providing appropriate fields of fter those operands to MULT 20314 and EXP 20316 in the formats 10 most efficiently utilized by MULT 20314 and EXP 20316.

I. In doing so, OPB 20322 extracts exponent and mantissa *i t Sfields from single and double precision floating point S operands to provide exponent and mantissa fields of these t" operands to, respectively, EXP 20316 and MULT 20314, and also unpacks, or converts, unpacked decimal operands to i packed decimal operands most efficiently utilized by MULT 20314.

Having described structure and operation of OPB VCi 20322, structure and operation of MULT 20314 will be described next below.

3. Multiplier 2Q314-(Pics-..257,- 258).

MULT 20314, as previously described, performs addition, subtraction, multiplication, and division operations on mantissa fields of single and double precision floating point operands, integer operands, and decimal operands. As described above with reference to OPB 20322, OPB 20322 converts unpacked decimal operands to packed decimal operands to be operated upon by MULT h St- 1 i g l I"-L 1^ ^T .r I -423- 20314. MULT 20314 is thereby effectively capable of performing all arithmetic operations on unpacked decimal operands.

a.a. Multiplier 20314 Data-Paths-and Memory fFia 257-) Referring to Fig. 257, a more detailed block diagram of MULT 20314's data paths and memory is shown. As previously described, major elements of MULT 20314 include memory elements comprised of MULTRF 20350 and S: 10 CONST 20360, operand input and result output multiplexing logic including MULTIM 20328 and MULTRM 20334, and t arithmetic operation logic. MULT 20314's operand input and result output multiplexing logic and memory elements Ft, will be described first, followed by description of MULT 20314's arithmetic operation logic.

As previously described, input data, including operands, is provided to MULT 20314's arithmetic Soperation logic through MULTIN Bus 20354. MULTIN Bus 20354 may be provided with data from three sources. A first source is CONST 20360 which is a 512 word by 32 bit wide Read Only Memory. CONST 20360 is utilized td store constants used in arithmetic operations. In particular, CONST 20360 stores zone fields for unpacked decimal, that is ASCI character, operands. As previously described, unpacked decimal operands are received by OPB 20322 and converted to packed decimal operands for more efficient utilization by MULT 20314. As such, final result outputs generated by MULT 20314 from such operands are in packed decimal format. As will be described below, MULT 20314 -424may be utilized to convert these packed decimal results into unpacked decimal results by insertion of zone fields. As indicated in Fig. 257, address inputs are provided to CONST 20360 from EXPQ Bus 20325 and from S output of mCRD 20346. Selection between these address inputs is provided through CONST Address Multiplexer (CONSTAM) 25710. CONST 20360 addresses will, in general, be provided from EUCL 20310 but alternately may be provided from EXPQ Bus 20325 for special operations.

10 Operand data is provided to MULTIN Bus 20354 through MULTIM 20328, which is a dual input, 64 bit multiplexer.

A first input of MULTIM 20328 is provided from OPQ Bus 20323 and is comprised of operand information provided from OPB 20322. OPQ Bus 20323 is a 56 bit wide bus and operand data appearing thereon may be comprised of 32 bit C 4 integer operands; 32 bit packed decimal operands, either C provided directly from OPB 20322 or as a result of OPB 20322's conversion of an unpacked decimal to a packed s, decimal operand; 24 bit single precision operand mantissa fields; or 56 bit double precision floating point operand mantissa fields. As previously described, certain OPQ Bus 20323 may be zero or sign extension filled, depending I' upon the particular operand.

Second input of MULTIM 20328 is provided from MULTRF 20350. MULTRF 20350 is a 16 word by 64 bit wide random access memory. As indicated in Figs. 203 and 257, MULTRF 20350 is connected between output of RFR 20336, through FR Bus 20337, and to input of MULT 20314's arithmetic operation logic through MULTIM 20328 and MULTIN Bus -425- 20354. MULTRF 20350 may therefore be utilized as a scratch pad memory for storing intermediate results of arithemetic operations, including reiterative arithmetic operations. In addition, a portion of MULTRF 20350 is utilized, as in GRF 10354, as an EU 10122 Stack Mechanism similar to MIS 10368 and MOS 10370 in FU 10120.

Operation of EU 10122 Stack Mechanism will be described in a following descripion of EU 10122's interfaces to MEM 10112 and FU 10120. Address Inputs (ADR) of MULTRF 20350 10 are provided from Multiplier Register File Address Multiplexer (MULTRFAM) 25712.

SMULTRFAM 25712 is a dual four bit multiplexer Sr comprised, for example, of SN74S258s. In addition to t address inputs to MULTRF 20350, MULTRFAM 25712 provides address inputs to EXPRF 20380. As previously described, MULTRF 20350 and EXPRF 20380 together comprise an EU Sn 10122 general register file similar to GRF 10354 and FU 10120. As such, MULTRF 20350 and EXPRF 20380 are CF addressed in parallel to read and write parallel entries from and to MULTRF 20350 and EXPRF 20380. Address inputs to MULTRFAM 25712 are provided, first, from outputs of mCRD 20346, thus providing microinstruction control of rcet addressing of MULTRF 20350 and EXPRF 20380. Second address input to MULTRFAM 25712 is provided from output of Multiplier Register File Address Counter (MULTRFAC) 25714.

MULTRFAC 25714 is a four bit counter and is used to generate sequential addresses to MULTRF 20350 and EXPRF 20380. Initial addresses are loaded into MULTRFAC 25714 -426from Multiplier Register File Address Counter Multiplexer (MULTRFACM) 25716. MULTRFACM 25716 is a dual four bit multiplexer. Inputs to MULTRFACM 25716 are provided, first, from outputs of mCRD 20346. This input allows microinstruction selection of an initial address to be loaded into MULTRFAC 25714 to be subsequently used and generating sequential MULTRF 20350 and EXPRF 20380 addresses. Second address input to MULTRFACM 25716 is provided from OPQ Bus 20323. MULTRFACM 25716's input 10 from OPQ Bus 20323 allows a single address, or a starting address of a sequence of addresses, to be selected through JPD Bus 10142 or MOD Bus 10144, for example from MEM 10112 or FU 10120.

Intermediate and final result outputs of MULT 20314 arithmetic logic are provided to data inputs of MULTRF 20350 directly from FR Bus 20337 and from MULTRM 20334.

Inputs to MULTRM 20334, in turn, are provided from FR Bus 20337 and from output of CONSIZE 20352 and TSTINT 20320.

'FR Bus 20337 is a 64 bit bus connected from 64 bit output of RFR 20336 and carries final and intermediate results of MULT 20314 arithmetic operations. As will t' become apparent in a following description of MULT 20314 Stt arithmetic operation logic, RFR 20336 output, and thus FR Bus 20337, are 64 bits wide. Sixty-four bits are provided to insure retention of all significant data bits of certain MULT 20314 arithmetic operation intermediate results, in p-rticular operations involving double precision floating point 64 bit mantissa fields. In addition, as will be described momentarily and has been L

A

'I 1 -427previously stated, MULT 20314 may convert a final result in packed decimal format into a final result in unpacked decimal format. In this operation, a single 32 bit, or one word, packed decimal result is converted into a 64 bit, or two word, unpacked decimal format by insertion of zone fields.

As described above, two parallel data paths are provided to transfer information from FR Bus 20337 into ,.lr MULTRF 20350. First path is directly from FR Bus 20337 10 and second path is through Unpacked Decimal Multiplexer (UPDM) 25718 of MULTRM 20334. Direct path is utilized for thirty-two bits of information comprising bits 0 to 23 and bits 56 to 63 of FR Bus 20337. Data path through UPDM 25718 may comprise either bits 24 to 55 of FR Bus 20337, which are connected into a first input of UPDM 25718, or bits 40 through 55 which are connected to a second input of UPDM 25718. Single precision floating point numbers are 32 bit numbers plus two or more guard bits and are thus written into MULTRF 20350 through bits 0 to 23 of the direct path into MULTRF 20350 and through first input (bits 24 to 55) of UPDM 25718. Double precision floating point numbers are 5 bits wide, plus guard bits, and thus utilize the direct path into MULTRF 20350 and the path through first input of UPDM 25718.

Bits 56 to 63 of direct path are utilized for guard bits of double precision floating point numbers. Both integer F and packed decimal numbers utilize bits 24 through 55 of FR Bus 20337, and are thus written into MULTRF 20350 through first input of UPDM 25718. As previously -428described, bits 0 to 23 of these operands are filled by sign extension.

a.a.a. Containe-r-Size--Check As stated above, MULTRM 20334 has an input from CONSIZE 20352. As will be described below with reference to TSTINT 20320, CONSIZE 20352 performs a "container size" check upon each store back of results from EU 10122 to MEM 10112. CONSIZE 20352 compares the number of significant bits in a result to be stored back to the 1, 0 logical descriptor describing the MEM 10112 address space that result is to be written into. Where reiterative write operations to MEM 10112 are required to transfer a result into MEM 10112, that is a string transfer, container size information may read from CONSIZE 20352 through Container Size Driver (CONSIZED) 25720 and MULTRM 20334 and written into MULTRF 20350. This allows EU 10122, using container size information stored in MULTRM 20350, to perform continuous container size checking during a string transfer of result from EU 10122 to MEM 10112. In addition, as will be described momentarily, container size information may be read from CONSIZE 20352 to JPD Bus 10144.

b.b.b. Final Result-Output Multiplexer-2Q0324 Referring finally to FROM 20324, as previously described FROM 20324 is utilized to transfer, in general, results of EU 10122 arithmetic operations onto JPD Bus 10142 for transfer to MEM 10112 or FU 10120. As indicated in Fig, 257, FROM 20324 is comprised of 24 bit Final Result Bus Driver (FRBD) 25722 and Result Bus -429- Driver (RBR) 25724. Input of FRBD 25722 is connected from FR Bus 20337 and allows data appearing thereon to be transferred onto JPD Bus 10142. In particular, FRBD 25722 is utilized to transfer 24 bit mantissa fields of single precision floating point results onto JPD Bus 10142 in parallel with a corresponding exponent field from EXP 20316. RBR 25724 input is connected from RSLT Bus 20388 to allow output of UPDM 25718 to be transferred onto JPD Bus 10142. RBR 25724, RSLT Bus 20388, and UPDM 10 25718 are used, in general, to transfer final results of EU 10122 operations from output of MULT 20314 onto JPD Bus 10142. Final results transferred by this data path rfct* include integer, packed and unpacked decimal results, and mantissa fields of double precision floating point results. Both unpacked decimal numbers and mantissa fields of double precision, floating point numbers are comprised of two 32 bit words and are thus transferred onto JPD Bus 10142 in two sequential transfer operations.

Having described structure and operation of MULT 20314's memory elements and input and output circuitry, MULT 20314's arithmetic operation logic will be described 9 9: next below.

4. Test- and- Int e+ .260-268) As previously described, TSTINT 20320 includes CONSIZE 20352, ECPT 20328, TSTCOND 20384, and INTRPT 20388. CONSIZE 20352, as previously described, performs "container size" check operations when results of EU -430- 10122 operations are to be written into MEM 10112. That is, CONSIZE 20352 compares size or number of significant bits, of an EU 10122 result to the capacity, or container size, of the MEM 10112 location that EU 10122 result is to be written into. As indicated, in Fig. 203, CONSIZE 20352 receives a first input, that is the results of EU 10122 operations, from FR Bus 20337. A second input of CONSIZE 20352 is connected to LENGTH Bus 20226 to receive length field of logical descriptors identifying MEM 10112 address space into which those EU 10122 results are to be written. CONSIZE 20352 includes logic circuitry, for example a combination of Read Only Memory arid Field Programmable Logic Arrays, for examining EU 10122 operation results appearing on FR Bus 20337 and determining the number of bits of data in those results.

CONSIZE 20352 compares EU 10122 result size to logical descriptor length field and, in particular, if result size exceeds logical descriptor length, provides an alarm output to ECPT 20328, described below.

TSTCOND 20384, previously described and which will be described further below, is an interface circuit between FU 10120 and EU 10122. TSTCOND 20384 allows FU 10120 to specify and examine results of certain test operations performed by EU 10122 with respect to EU 10122 operations.

ECPT 20328 monitors certain EU 10122 operations and provides outputs indicating when certain "exceptions" have occurred. These exceptions include attempted 1 I- J 1 -431divisions by zero, floating point exponent underflow or overflow, and integer container size fault.

INTRPT 20388 is again an interface between EU 10122 and FU 10120 allowing FU 10120 to interrupt EU 10122 operations. INTRPT 20388 allows FU 10120 to direct EU 10122 to execute certain operations to aid in handling of certain FU 10120 events previously described.

Operation of CONSIZE 20352, ECPT 20328, TSTCOND 20384, INTRPT 20388, and other features of EU 10122's S, 10 interface with FU 10120 will be described further below t f in the following description of operation of that interface and of operation of certain EU 10122 internal mechanisms, such as FU 10120 Stack Mechanisms.

a.a. -FU120aEU-1012-2Interiace As previously described, EU 10122 and FU 10120 are asychronous processors, each operating under its own Smicrocode control. EU 10122 and FU 10120 operate simultaneously and independently of each other but are coupled, and their operations coordinated, by interface signals described below. Should EU 10122 not be able to respond immediately to a request from FU 10120, FU 10120 will idle until EU 10122 becomes available; conversely, should EU 10122 not receive, or have present, operands or a request for operations from FU 10120, EU 10122 will remain in idle state until operands and requests for operations are received from FU 10120.

In normal operation, EU 10122 manipulates operands under control of FU 10120, which in turn is under control i -432of SOPs of a user's program. When FU 10120 requires arithmetic or logical manipulation of an operand, FU 10120 dispatches a command, that is an Execute Unit Dispatch Pointer (EUDP) to EU 10122. As previously described, an EUDP is basically an initial address into EOSITT 20344. An EUDP identifies starting location of a EU 10122 microinstruction sequence performing the required operation upon operands. Operands are fetched from MEM 10112 under FU 10120 control, as previously described, and are transferred into OPB 20322. Those operands are then called from OPB 20322 by EU 10122 and transferred into MULT 20314 and EXP 20316 as previously described. After the required operation is completed, FU 10120 is notified that a result is ready. At this point, FU 10120 may check certain test conditions, for example through TSTCOND 20384, such as whether an integer or decimal carry bit is set or whether a mantissa sign bit t is set or reset. This test operation is utilized by FU 10120 for conditional branching and synchronization of FU 10120 and EU 10122 operations,' Exception checking, by ECPT 20328, is also performed at this time. Exception checking determines, for example, whether division by zero was attempted or if a container size fault has occurred. In general, FU 10120 is not informed of exception errors until FU 10120 requests exception checking. After results are transferred into FU 10120 or MEM 10112 by EU 10122, EU 10122 goes to idle operation i until a next opcra~ion is requested by FU 10120.

i L. 1 i v -433- Having briefly described overall interface operation between FU 10120 and EU 10122, operation of that interface, referred to as handshaking, will be described in greater detail next below. In general, handshaking operation between EU 10122 and FU 10120 during normal operation may be regarded as following into six operations. These operations may include, for example, loading of COMQ 20342, loading of OPB 20322, storeback or transfer of results from EU 10122 to FU 10120 or MEM S. 10 10112, check of test conditions, exception checking, and S. EU 10122 idle operation. Handshaking between FU 10120 and EU 10122 will be described below for each of these 1 JAclasses of operation, in the order just referred to.

aa.aa. Loadina-of--Command Ouuee-2034-2- (Pi Referring to Fig. 260, a schematic representation of EU 10122's interface with FU 10120 for purposes of loading COMQ 20342 as shown. During normal SOP directed JP 10114 operation, 8 bit operation (OP) codes are parsed Ifl from the instruction stream, as previously described, and concatenated with dialect information to address EUSDT 20266 also as previously described. EUSDT 20266 provides corresponding addresses, that is EUDPs, to EUSITT 20344.

Dialect information specifies the S-Language currently being executed and, consequently, the group of microinstruction sequences available in EUSITT 20344 for that S-Language. As previously described, FU 10120 may specify four S-Language dialects with up to 256 EU 10122 i iI. -434microinstruction sequences per dialect, or 8 dialects with up to 128 microinstruction sequences per dialect.

EUDPs provided by EUSDT 20266 are comprised of a 9 bit address field, a 2 bit operand information field, and a 1 bit flag field, as previously described. Address field is starting address of a microinstruction sequence ifEUSITT 20344 and EU 10122 will perform the operation S'

J

directed by that microinstruction sequence. EUSITT 20344 requires 11 bits of address field and the 9 bit address field of EUDPs are mapped into an 11 bit address field by left justification and zero filling.

SFU 10120 may also dispatch, or select, any EU 10122 t microinstruction controlled operation from JPD Bus 10142.

Such EUDPs are provided from JPD Bus 10142 to data input of EUSITT 20344 and passed directly through to mCRD *r 20346. Before a EUDP may be provided from JPD Bus 10142, however, FU 10120 provides a check operation comparing that EUDP to a list of legal, or allowed, EUSITT 20344 addresses stored in MEM 10112. A fault will be indicated if an EUDP provided through JPD Bus 10142 is not a legal S EUSITT 20344 address. Alternately, FU 10120 may effectively provide an EUDP, or EUSITT 20344 addresses, from a literal field in a FU 10120 microinstruction word. Such a FU 10120 microinstruction word literal field may be effectively utilized as an SOP into EUSDT 20266.

Handshaking between EU 10122 and FU 10120 during load COMQ 20342 operations may proceed as illustrated in Fig.

1 4' -435- 260. A twelve bit EUDP may be placed on EUDIS Bus 20206 and Control Signal Load Command Queue (LDCMQ) asserted.

If COMQ 20342 is full, EU 10122 raises control signal Command Hold (CMDHOLD) which causes FU 10120 to remain in State MO until there is room in COMQ 20342. As previously described, COMQ 20342 is comprised of two, two word buffers wherein one buffer is utilized for normal SOP operation and the other utilized for control of FU 10120 and EU 10122 internal mechanism operation.

S 10 EUDPs are loaded into COMQ 20342 when state timing signals M1CPT and M1 are asserted. If a EUDP being transferred into COMQ 20342 concerns a double precision C cc floating point operation, control signal Set Double Precision (SETDP) is asserted. SETDP is utilized to control OPB 20322, and because single precision and precision floating point operations otherwise utilize the same SOP and thus would otherwise refer to same EUSITT 20344 microinstruction sequence.

t At this point, a EUDP has been loaded into COMQ 20342 and will be decoded to control FU 10120 operation by EUCL 20310 as previously described. Each particular EUDP will be cleared by that EUDPs EUSITT 20344 microinstruction S. t t sequence after the requested microinstruction sequence has been executed.

b.b.b. Loading"-of-Operand Buffer'.2 3O2o--i Pig 61' Referring to Fig. 261, a diagramic representation of the interface and handshaking between EU 10122, FU 10120 and MEM 10112 for loading OPB 20322 is shown. Control signal Clear Queue Full (CLQF) from EU 10122 must be ii f *i -436r4,9 9 99 .9 4 I 4t 4 4 ra I 4(e4 4 4 asserted by EU 10122 before FU 10120 initiates a request to MEM 10112 for an operand to be transferred to EU 10122. CLQF clears and "EU 10122's OPB 20322 Full" condition in FU 10120. CLQF indicates, thereby, that there is room in OPB 20322 to receive operands. If FU 10120 is in a "EU 10122's OPB 20322 Full" condition and further operand is required to be transferred to EU 10122, FU 10120 will remain in State Ml until CLQF is asserted.

10 At the beginning of execution of a particular SOP, FU 10120 may transfer two operands to OPB 20322 without "EU 10122's OPB 20322 Full" condition occurring. This is because EU 10122 is idle at the beginning of an SOP execution and generally immediately unloads a first operand from OPB 20322 before a second operand arrives.

Control signal Job Processor Operand (JPOP) provided from FU 10120 must be non-asserted for operands to be transferred from MEM 10112 to OPB 20322 through MOD Bus 10144. This is the normal condition of JPOP. If JPOP is 20 asserted, OPB 20322 is loaded with data from JPD Bus 10142. Data is strobed into OPB 20322 from JPD Bus 10142 by control signals MICPT and JPOP. Operands read from MEM 10112, however, are transferred into OPB 20322 through MOD Bus 10144 when MEM 10112 asserts DAVEB to indicate that valid data from MEM 10112 is available on MOD Bus 10144. DAVEB is also utilized to strobe data on MOD Bus 10144 into OPB 20322. If control signal ZFILL from MEM 10112 is asserted at this point, ZFILL is 1iU

I~

i i;' %i ~if~i~i c i- t >4 -437interpreted during integer operand operations to indicate that those operands are unsigned and should be left zero filled, rather than sign extended. If data is being provided from JPD Bus 10142 rather than from MEM 10112, that is if JPOP is asserted, bit 11 of current EUDP may be utilized to perform the same function as ZFILL during loading of OPB 20322 from MOD Bus 10144.

Loading of OPB 20322 is controlled, in part, by bits .I 9 and 10 of EUDPs provided from FU 10120 through BUDIS I 10 Bus 20206. Bit 9 indicates length of a first operand S• while bit 10 indicates length of a second operand.

C Operand length, together with operand type specified in e address portion of a EUDP, determines how a particular operand is unloaded from OPB 20322 and transferred into MULT 20314 and EXP 20316.

At this point, both. COMQ 20342 and OPB 20322 have been loaded with, respectively, EUDPs and operands. It should be noted that operands are generally not r \transferred into OPB 20322 before a corresponding EUDP is loaded into COMQ 20342. Operands and EUDPs may, however, be simultaneously transferred into EU 10122. If other operands are required for a particular operation, those operands are loaded into OPB 20322 as described above.

c.c.c. Stor eback-ei-; -2 6-2 Referring to Fig. 262, a diagramic representation of a storeback, or transfer, of results to MEM 10112 from EU 10122 and handshaking performed therein is shown. When a Sfinal result of a EU 10122 operation is available, EU 10122 asserts control signal Data Ready (DRDY). FU 10120 thereupon responds with control signal Transfer to JPD J41&'w r -438- Bus 10142 (XJPD), which gates EU 10122's result onto JPD Bus 10142. In normal operation, that is execution of SOPs, FU 10120 causes EU 10122's result to be stored back into a destination in MEM 10112, as selected by a physical descriptor provided from FU 10120. Alternately, a result may be transferred into FU 10120, 32 bits, or one word, at a time.

FU 10120 may, as described above and described further below, check EU 10122 test conditions during 10 storeback of results. FU 10120 generates control signal tI:' Transfer Complete (XFRC) once the storeback operation is completed. XFRC also indicates to EU 10122 that EU i 10122's results. and test conditions have been accepted by FU 10120, so that EU 10122 need no longer assert these results and test conditions.

j t d.d.d. TPest-~-ond_ tns--_-F Referring to Fig. 263, a diagramic representation of checking of EU 10122 test conditions by FU 10120, and handshaking therein, is shown. As previously described, test results indicating certain conditions and operations of EU 10122 are sampled and stored in TSTCOND 20384 and may be examined by FU 10120. When DRDY is asserted by EU 10122, FU 10120 may select, for example, one of 8 EU 10122 conditions to test, as well as transferring results as described above. EU 10122 conditions which may be tested by FU 10120 are listed and described below. Such conditions, as whether a final result is positive, -439negative, or zero, may be checked in order to facilitate conditional branching of FU 10120 operations as previously described. FU 10120 specifies a condition to be tested through Test Condition Select signals (TEST(2- FU 10120 asserts control signal EU Test Enable (EUTESTEN) to EU 10122 to gate the selected test condition. That selected test condition then appears as Data Signal Test Condition (TC) from EU 10122 to FU I c 10120. A TC of logic 1 may, for example, indicate that 10 the selected condition is false while a TC of logic 0 may S.indicate that the selected condition is true. FU 10120 C indicates that FU 10120 has sensed the requested test S" condition, and that the test condition need no longer be asserted by EU 10122, by asserting control signal XFRC.

e.e.e. Exception-ehecking gReferring to Fig. 264, a diagramic representation of exception checking of EU 10122 exceptions by FU 10120, and handshaking therein, is shown. As previously described, any EU 10122 exception conditions may be checked by FU 10120 as FU 10120 is initiating storeback of EU 10122 results. Exception checking may detect, for example, attempted division by zero, floating point "exponent underflow or overflow, or a container size fault. An attempted division by zero or floating point underflow or overflow may be checked before storeback, that is without specific request by FU 10120.

As previously described, a container size fault is detected by CONSIZE 20352 by comparing length of result with size of destination container in MEM 10112.

A-"

-440- Container size exception checking occurs during store back of EU 10122 results, that is while FU 10120 is in State SB. Container size is automatically performed by EU 10122 hardware, that is by CONSIZE 20352, only on results of less than 33 bits length. Size checking of larger results, that is larger integers and BCD results, is performed by a microcode routine, using CONSIZE 20352's output, as transfer of such larger results is executed as string transfer. It is unnecessary to 10 perform container size check for either single or double H precision floating point results as these data types L t t always occupy either 32 or 64 bits. Destination S' container size is provided to CONSIZE 20352 through c LENGTH Bus 20226.

Control signal Length to Memory AON or Random Signals C (LMAONRS) is generated by FU 10120 from Type field of the Slogical descriptor corresponding to a particular EU 10122 result. LMAONRS indicates that the results data type is J van unsigned integer. LMAONRS determines the manner in which a required container size of the EU 10122 result is determined. After receivng this information from LMAONRS, EU 10122 determines whether destination container size in MEM 10112 is sufficiently large to contain the EU 10122 result. If that destination container size is not sufficiently large, a container size fault is detected, by CONSIZE 20352, or through an EU 10122 microinstruction sequence.

i

.)J

-441-

I

t 4 t I 1 (4 It I 1 41 1C iI 'I I 1-11 1

I

$11111 I I

CC

tc- 44r 4, 4

I

I I I I It c, 1 (I I 4 44 44 4 4 4 41 I 4 4 Container size faults, as well as division by zero and exponent underflow and overflow faults, are signaled to FU 10120 when FU 10120 asserts control signal Check Size (CKSIZE). At this time, EU 10122 asserts control signal Exception (EXCPT) if any of the above faults has occurred. If a fault has occurred, an Event request to FU 10120 results. When an Event request is honored by FU 10120, FU 10120 may interrupt EU 10122 and dispatch EU 10122 to a microinstruction routine that transfers those 10 exception conditions onto JPD Bus 10142. If a container size fault has caused that exception condition, EU 10122 may transfer to FU 10120 the required container size through JPD Bus 10142.

f.f.f. -Idle.Routine 15 Finally, when a current EU 10122 operation is completed, EU 10122 goes into an Idle loop microinstruction routine. If necessary, FU 10120 may assert control signal Excute Unit Abort (EUABORT) to force EU 10122 into Idle loop microinstruction routine until EU 10122 is required for further operations.

g.g.g. EU..10,422-Stack Mechaniams..4-asg--265-.

As previously described, EU 10122 may perform either of two classes of operations. First, EU 10122 may perform arithmetic operations in execution of SOPs of user's programs. Second, EU 10122 may operate as an arithmetic calculator assisting operation of FU 10120's ii 4'~

\C

N>~

1 ff, 1 -442internal mechanisms and operations, referred to as kernel operations.

In kernel operation, EU 10122 acts as an arithmetic calculator for FU 10120 during address generation, address translation, and other kernel functions. In kernel mode, EU 10122 is executing microinstruction sequences at request of FU 10120 kernel microinstruction sequences, rather than at request of an SOP. In general, S these kernel operations-are vital to operation of JP 10 10114. FU 10120 may interrupt EU 10122 operations with regard to SOPs and initiate EU 10122 microinstruction sequences to perform kernel operations.

I When interrupted, EU 10122 saves EU 10122's current l operating state in a one level deep stack. EU 10122 may then accept an EUDP from that portion of COMQ 20342 I utilized to receive and store EUDPs regarding FU 10120's Iand EU 10122's internal, or kernel, operations. When requesting kernel operations by EU 10122, FU 10120 V generally transfers operands to OPB 20322 through JPD Bus 10142, and receives EU 10122 final results through JPD Bus 10142. Operands may also be provided to EU 10122 through MOD Bus 10144. After EU 10122 has completed a *i requested kernel operation, EU 10122 reloads operating j state from its internal stack and continues normal operation from the point normal operation was interrupted. 1 Should another interrupt from FU 10120 occur while a l prior interrupt is being executed, EU 10122 moves current r t L 1 -443state and data, that is of first interrupt, to MEM 10112. EU 10122 requests FU 10120 store state and date of first interrupt in MEM 10112 by requesting an "EU 10122 Stack Overflow" Event. EU 10122's "normal" state, that is state and data pertaining to the operation EU 10122 is executing at time of occurrence of first interrupt, is stored in an EU 10122 internal stack and remains there. EU 10122 then begins executing second interrupt. When EU 10122 has completed operations for 10 second interrupt, state from first interrupt is reloaded from MEM 10112 by EU 10122 requesting a "EU 10122 Stack Underflow" Event to FU 10120. EU 10122 then completes t execution of first interrupt and reloads state and resumes execution of normal operation, that is the operation being executed before the first interrupt.

CEU 10122 is therefore capable of handling interrupts from FU 10120 during two circumstances. First interrupt circumstance is comprised of interrupts occurring during normal operation, that is while executing SOPs of user's programs. Second circumstance arises when interrupts occur during kernel operations, that is during execution of microinstruction sequences for handling interrupts.

EU 10122 operation will be described next below for each of these circumstances, and in the order referred to.

Referring to Fig. 265, a diagramic representation of EU 10122's stack mechanisms, previously described, is shown. Those portions of EU 10122's stack mechanisms residing within EU 10122 aze comprised of EU 10122's Current State Registers (EUCSRs) 26510 and EU 10122's

I

-444- Internal Stack (EUIS) 26512. EUCSR 26510 is comprised of EU 10122's internal registers which contain data and state of current EU 10122 operation. EUCSR 26510 may be comprised, for example, of mCRD 20346, .registers of TSTINT 20320, and the previously described registers within MULT 20314 and EXP 20316.

State and data contained in EUCSR 26510 is that of the operation currently being executed by EU 10122. This current state may, for example, be that of a SOP 10 currently being executed by EU 10122, or that of an interrupt, for example a fourth interrupt of a nested sequence of interrupts, requested by FU 10120.

EUIS 26512 is comprised of certain registers of MULTRF 20350 and EXPRF 20380. EUIS 26512 is utilized to store and save current state of an SOP operation Scurrently being executed by EU 10122 and which has been interrupted. State and data of that SOP operation will remain stored in EUIS 26512 regardless of the number of interrupts which may occur on a nested sequence of t 20 interrupts requested by FU 10120. State and data of the interrupted SOP operation will be returned from EUIS 26512 to EUCSR 26510 when all interrupts have been Scompleted.

Final portion of EU 10122's stack mechanism is that portion of EU 10122's internal stack (EUES) 26514 residing in MEM 10112. EUES 26514 is comprised of certain MEM 10112 address locations used to store state and data of successive interrupt operations of sequences of nested interrupts. That iso if a sequence of four -445interrupts is requested by FU 10120, state and data of fourth interrupt will reside in EUCSR 26510 while state and data of first, second, and third interrupts have been transferred, in sequence, into EUES 26514. In this respect, and as previously described operation of EU 10122's stack mechanisms is similar to that of, for example, MIS 10368 and SS 10336 previously described with reference to Fig. 103., As described above, an interrupt may be requested of 10 EU 10122 by FU 10120 either during EU 10122 normal operation, that is during execution of SOPs by EU 10122, or while EU 10122 is executing a previous interrupt ctt requested by FU 10120. Operation of EU 10122 and FU 10120 upon occurrence of an interrupt during EU 10122 normal operation will be described next below.

Referring to Fig. 266, a diagramic representation of handshaking between EU 10122 and FU 10120 during an interrupt of EU 10122 while EU 10122 is operating in normal mode is shown and should be referred to in conjunction with rig. 265. For purposes of the following discussions, interrupts of EU 10122 operations by FU 10120 are referred to as nanointerrupts to distinguish lrf from interrupts internal to FU 10120.

FU 10120 interrupts normal operation of EU 10122 by assertion of control signal Nano-Interrupt (NINTP) during State MO of FU 10120 operation. NINTP may be masked by EU 10122 during certain critical EU 10122 operations, such as arithmetic operations. If NINTP is masked by EU 10122, FU 10120 will remain in State NW until EU 10122 acknowledges the interrupt.

L

__72~ -446- Upon receiving NINTP from FU 10120, EU 10122s transfers state and data of current SOP operation from EUCSR 26510 to EUIS 26512. EU 10122 then asserts control signal Nano-Interrupt Acknowledge (NIACK) to FU 10120 to acknowledge availability of EU 10122 to accept a nanointerrupt. FU 10120 will then enter State Ml and place an EUDP on EUDIS Bus 20206. Loading of COMQ 20342 then proceeds as previously described, with EU 10122 I, ,t loading nanointerrupt EUDPs into the appropriate 10 registers of COMQ 20342. COMQ 20342 is loaded as previously described and, if JPOP is asserted, data transferred into OPB 20322 from JPD Bus 10142. If JPOP is not asserted, data is taken into OPB 20322 fro MOD Bus 10144. EU 10122 then proceeds to execute the required nanointerrupt operation and storing back of results and checking of test conditions proceeds as previously described for EU 10122 normal operation. In general, exception checking is not performed. When EU t 10122 has completed execution of the nanointerrupt operation, EU 10122 transfers state and data of the interrupted SOP operation from EUIS 26512 to EUCSR 26510 S. and resumes execution of that SOP. At this point, EU I10122 asserts control signal Nano-Interrupt Trap Enable (NITE). NITE is received and tested by FU 10120 to indicate end of nanointerrupt processing.

Referring to Fig. 267, a diagramic representation of interfaces between EU 10122, FU 10120, and MEM 10112 during nested, or sequential, EU 10122 interrupts for -447kernel operations, and handshaking therein, is shown.

During the following discussion, it is assumed that EU 10122 is already processing a nanointerrupt for a kernel operation submitted to EU 10122 by FU 10120. FU 10120 may then submit a second, third, or fourth, nanointerrupt to EU 10122 for a further kernel operation. FU 10120 will assert NINTP to request a nanointerrupt of EU 10122. EU 10122's normal mode state and data from a 1 previously executing SOP operation has been stored in, S and remains in, EUIS 26512. Current state and data of Scurrently executing nanointerrupt operation in EUCSR S26510 will be transferred to EUES 26514 in MEM 10112 to allow initiation of pending nanointerrupt. EU 10122 will at this time assert NIACK and control signal Execute Unit Event (EXEVT). EXEVT to FU 10120 informs FU 10120 that an EU 10122 Event has occurred, specifically, and in this case, EXEVT requests FU 10120 service of an EU 10122 Stack Overflow. FU 10120 is thereby trapped to an "EU S10122 Stack Overflow" Event Handler microinstruction sequence. This handler transfers current state and data of interrupted nanointerrupt previously executing in EU t (r St 10122 into EUES 26514. State and data of interrupted ce nanointerrupt is transferred to EUES 26514, one 32 bit word at a time. FU 10120 asserts control signals XJPD to gate each of these state and data words onto JPD Bus 10142 and controls transfer of these words into EUES 26514.

Processing of new nanointerrupt proceeds as described above with reference to interrupts occurring during L 2 4 Prcesingof ew nnoiteruptprocedsas escibe -448normal operation. If any subsequent nanointerrupts occur, they are handled in the same manner as just Cescribed; FU 10120 signals a nanointerrupt to FU 10120, current EU 10122 state and data is saved by FU 10120 in EUES 26514, and new nanointerrupt is processed. After a nested nanointerrupt, that is a nanointerrupt of a sequence of nanointerrupts, has been serviced, EU 10122 asserts control signal EU 10122 Trap (ETRAP) to FU 10120 to request a transfer of a previous nanointerrupt's state S 10 and data from EUES 26514 to EUCSR 26510. FU 10120 will t ^retrieve that next previous nanointerrupt state and data t from EUES 26514 through MOD'Bus 10144 and will transfer that data and state onto JPD Bus 10142. This state and data is returned, one 32 bit word at a time, and is strobed into EU 10122 by JPOP from FU 10120. Processing of that prior nanointerrupt will then resume. The servicing of successively prior nanointerrupts will Scontinue until all previous nanointerrupts have been serviced. Original state and data of EU 10122, that is that of SOP operation which was initially interrupted, is then returned to EUCSR 26510 from EUIS 26512 and execution of that SOP resumed. At this time, EU 10122 asserts NITE to indicate end of EU 10122 kernel operations in regard to nanointerrupts.

Having described structure and operation of EU 10122, FU 10120 and MEM 10112, with respect to servicing of kernel operation nanointerrupts by EU 10122, loading of -449- EU 10122's EUSITT 20344 with microinstruction sequences will be described next below.

h.h.h. Loading--of--Execute--Unrit S-Interpreter--abe 20344--fP-ia---26 Referring to Fig. 268, a diagramic representation of interface and handshaking between EU 10122, FU 10120, MEM 10112, and DP 10118 during loading of microinstructions into EUSITT 20344 is shown. As previously described, EUSITT 20344 contains all microinstructions required for t control of EU 10122 in executing kernel nanointerrupt operations and in executing arithmetic operations in response to SOPs of user's programs. EUSITT 20344 may store microinstruction sequences for interpreting arithmetic SOPs of user's programs for, for example, up to 4 different S-Language Dialects. In general, a capacity of storing microinstruction sequences for arithmetic operations in up to 4 S-Language Dialects is sufficient for most requirements, so that EUSITT 20344 need be loaded with microinstruction sequences only at initialization of CS 10110 operation. Should t microinstruction sequences for arithmetic operations of more than 4 S-Language Dialects be required, those microinstruction sequences may be loaded into EUSITT 20344 in the manner as will be described below.

As previously described, a portion of the microinstructions stored in EUSITT 20344 is contained in Read Only Memories and is thus permanently stored in EUSITT 20344. Microinstruction sequences permanently 't -450stored in EUSITT 20344 are, in general, those required for execution of kernel operations. Microinstruction sequences permanently stored in EUSITT 20344 include those used to assist in writing other EU 10122 microinstruction sequences into.EUSITT 20344 as required. Certain microinstruction sequences are stored in a Random Access Memory, referred to as the Writeable Control Store (WCS) portion of EUSITT 20344, and include these for interpreting arithmetic operation SOPs of S 10 various S-Language Dialects.

c* Writing of microinstruction sequences into EU 10122 is initialized by forcing EU 10122 into an Idle state.

1 Initialization of EU 10122 is accomplished by FU 10120 asserting EUABORT or by DP 10118 asserting control signal clear (CLEAR). Either EUABORT or CLEAR will clear a current operation of EU 10122 and force EU 10122 into Idle state, wherein EU 10122 waits for further EUDPs provided from FU 10120. FU 10120 then dispatches a EUDP initiating loading of EUSITT 20344 to EU 10122 through EUDIS Bus 20206. Load EUSITT 20344 EUDP specifies a starting address of a two step microinstruction sequence in the PROM portion of EUSITT 20344. This two step microinstruction sequence first loads zeros into SCAG 25536, which as previously described provides read and write addresses to EUSITT 20344. EUSITT 20344 load microinstruction sequence then reads a microinstruction from EUSITT 20344 to mCRD 20346. This microinstruction specifies conditions for handshaking operations with FU 10120 so that loading of EUSITT 20344 may begin. At this 1 -451time, and from this microinstruction word, EU 10122 asserts control signal DRDY to FU 10120 to indicate that EU 10122 is ready to accept EUDPs from FU 10120 for directing loading of EUSITT 20344. This initial microinstruction also generates a write enable control signal for the WCS portion of EUSITT 20344, inhibits loading of mCRD 20346 from EUSITT 20344, and inhibits normal loading operations of NXTR 25540 and SCAG 25536.

ti" c This first microinstruction also directs NASS 25526 to f t 10 accept address inputs from SCAG 25536 and, finally, causes NITE to FU 10120 to be asserted to unmask nanointerrupts from FU 10120'.

S* FU 10120 then generates a read request to MEM 10112, and MEM 10112 transfers a first 32 bit word of a EU 10122 microinstruction word onto JPD Bus 10142. Each such 32 bit word from MEM 10112 comprises one half of a 64 bit microinstruction word of EU 10122. When FU 10120 receives DRDY from EU 10122, FU 10120 generates control II" signal Load Writeable Control Store (LDWCS). LDWCS in turn transfers a 32 bit word on JPD Bus 10142 into a first address of the WCS portion of EUSITT 20344. A next 32 bit half word of a EU 10122 microinstruction word is .then read from MEM 10112 through JPD Bus 10142 and transferred into the second half of that first address within the WCS portion of EUSITT 20344. The address in SCAG 25536 is thenincremented to select a next address within EUSITT 20344 and the process just described repeated automatically, including generation of DRDY and LDWCS, until loading of EUSITT 20344 is completed.

L

j i. y

-I

-452- After loading of EUSITT 20344 is completed, the loading process is terminated when FU 10120 asserts NINTP, or DP 10118 asserts Control Signal Load Complete (LOADCR). Either NINTP or LOADCR releases control of operation of NAG 20340 to allow EU 10122 to resume normal operation.

The above descriptions have described structure and operation of EU 10122, including: execution of various arithmetic operations utilizing various operand formats; S 10 operation of EU 10122, FU 10120, and MEM 10112 with Sregard to handshaking; loading of EUDPs and operands; storeback of results; checking of test conditions and exceptions; EU 10122 Stack Mechanisms during normal and kernel operations; and loading of EU 10122 microinstruction sequences into EUSITT 20344. IOS 10116 C fand DP 10118 will be described next below, in that order.

t t t s t D. -I/Q-0 S-stem- 3---sos---20s- Referring to Fig. 204, a partial block diagram of IOS 10116 is shown. As previously described, IOS 10116 operates as an interface between CS 10110 and the external world, for example, ED 10124. A primary function of IOS 10116 is the transfer of data between CS 10110, that is MEM 10112, and the external world. In 1 addition to performing transfers of data, IOS 10116 controls access between various data sources and sinks of ED 10124 and MEM 10112. As previously described, IOS 10116 directly addresses MEM 10112's physical address space to write data into or read data from MEM 10112. As such, IOS 10116 also performs address translation, a l I c- If -453mapping operation required in transferring data between MEM 10112's physical address space and address spaces of data sources and sinks in ED 10124.

As shown in Fig. 204, IOS 10116 includes Data Mover (DMOVR) 20410, Input/Output Control Processor (IOCP) 20412, and one or more data channel devices. IOS 10116's data channel devices may include ECLIPSE* Burst Multiplexer Channel (EBMC) 20414, NOVA Data Channel (NDC) .20416, and other data channel devices as required for a I. 10 particular configuration of a CS 10110 system. IOCP 20412 controls and directs transfer of data between MEM S" 10112 and ED 10124, and controls and directs mapping of r addresses between ED 10124 and MEM 10112's physical address space. IOCP 20412 may be comprised, for example, of a general purpose computer, such as an ECLIPSE* M600 computer available from Data General Corporation of Westboro, Massachusetts.

EBMC 20414 and NDC 20416 comprise data channels through which data is transferred between ED 10124 and IOS 10116. EBMC 20414 and NDC 20416 perform actual transfers of data to and from ED 10124, under control of IOCP 20412, and perform mapping of ED 10124 addresses to MEM 10112 physical addresses, also under control of IOCP 20412. EBMC 20414 and NDC 20416 may respectively be comprised, for example, of an ECLIPSE* Burst Multiplexer Data Channel and a NOVA" Data Channel, also available from Data General Corporation of Westboro, Massachusetts, DMOVR 20410 comprises IOS 10116's interface to MEM 10112. DMOVR 20410 is the path through which data and -454addresses are transferred between EBMC 20414 and NDC 20416 and MEM 10112. Additionally, DMOVR 20410 controls access between EBMC 20414, NDC 20416, and other IOS 10116 data channels, and MEM 10112.

ED 10124, as indicated in Fig. 204, may be comprised of one or more data sinks and sources. ED 10124 data sinks and sources may include commercially available disc drive units, line printers, communication lengths, tape units, and other computer systems, including other CS 10110 systems. In general, ED 10124 may include all such data devices as are generally interfaced with a computer 9 SI• system.

a. 9 Referring first to the overall structure of IOS 10116, data input/output of ECLIPSEO Burst Multiplexer it .Channel Adapter and Control Circuitry (BMCAC) 20418 of to EBMC 20414 is connected to bi-directional BMC Address and Data (BMCAD) Bus 20420. BMCAD Bus 20420 in turn is connected to data and address inputs and outputs of data sinks and sources of ED 10124.

Similarly, data and address inputs and outputs of NOVA* Data Channel Adapter Control Circuits (NDCAC) 20422 S"in NDC 20416 is connected to bi-directional NOVA* Data Channel Address and Data (NDCAD) Bus 20424. NDCAD Bus 20424 in turn is connected to address and data inputs and outputs of data sources and sinks of ED 10124. BMCAD Bus 20420 and NDCAD Bus 20424 are paths for transfer of data and addresses between data sinks and sources of ED 10124 °i -455and IOS 10116's data channels and may be expanded as required to include other IOS 10116 data channel devices and other data sink and source devices of ED 10124.

Within EBMC 20414, bi-directional data input and output of BMCAC 20418 is connected to bi-directional input and output of BMC Data Buffer (BMCDB) 20426. Data inputs and data outputs of BMCDB 20426 are connected to, 1 i respectively, Data Mover Output Data (DMOD) Bus 20428 and S. Data Mover Input Data (DMID) Bus 20430. Address outputs of BMCAC 20418 are connected to address inputs of Burst 4 Multiplexer Channel Address Translation Map (BMCATM) 20432 and address outputs of BMCATM 20432 are connected onto DMID Bus 20430. A bi-directional control input and output of BMCATM 20432 is connected from bi-directional 1O Control Processor Control (IOCPC) Bus 20434.

Referring to NDC 20416, as indicated in Fig. 204 data Sinputs and outputs of NDCAC 20422 are connected, respectively, from DMOD Bus 20428 and to DMID Bus 20430.

Address outputs of NDCAC 20422 are connected to address inputs of NOVA* Data Channel Address Translation Map (NDCATM) 20436. Address outputs of NDCATM 20436 are, in turn, connected onto DMID Bus 20430. A bi-directional control input and output of NDCATM 20436 is connected from IOCPC Bus 20434.

Referring to IOCP 20412, a bi-directional control input and output of IOCP 20412 is connected from IOCPC Bus 20434. Address and data output of IOCP 20412 is connected to NDCAD Bus 20424. An address output of IOCP Address Translation Map (IOCPATM) 20438 within IOCP 20412 4 -456is connected onto DMID Bus 20430. Data inputs and outputs of IOCP 20412 are connected, respectively, to DMOD Bus 20428 and DMID Bus 20430. A bi-directional control input and output of IOCP 20412 is connected to a bi-directional control input and output of DMOVR 20410.

Referring finally to DMOVR 20410, DMOVR 20410 includes Input Data Buffer (IDB) 20440, Output Data Buffer (ODB) 20442, and Priority Resolution and Control (PRC) 20444. A data and address input of IDB 20440 is 10 connected from DMID Bus 20430. A data and address output of IDB 20440 is connected to IOM Bus 10130 to MEM 10112.

A data output of ODB 20442 is connected from MIO Bus f" 10129 from MEM 10112, and a data output of ODB 20442 is connected to DMOD Bus 20428. Bi-directional control inputs and outputs of IDB 20440 and ODB 20442 are connected from bi-directional control inputs and outputs of PRC 20444. A bi-directional control input and output g, of PRC 20444 is connected from a bi-directional control input and output of IOCP 20412 as described above.

Another bi-directional control input and output of PRC 20444 is connected to and from IOMC Bus 10131 and thus ,1 from a control input and output of MEM 10112. Having described overall structure of IOS 10116, operation of IOS 10116 will be described next below.

b. *I/Q-System-10--Opera+on-+q-.- 9 1. Data-,nhanne--DPeva-ces Referring first to EBMC 20414, BMCAC 20418 receives data and addresses from ED 10124 through BMCAD Bus 17ly -457- I i« I 4 44 4 44 444 4

C

20420. BMCAC 20418 transfers data into BMCDB 20426, where that data is held for subsequent transmission to MEM 10112 through DMOVR 20410, as will be described below. BMCAC 20418 transfers addresses received from ED 10124 to BMCATM 20432. BMCATM 20432 contains address mapping information correlating ED 10124 addresses with MEM 10112 physical addresses. BMCATM 20432 thereby provides MEM 10112 physical addresses corresponding to ED 10124 addresses provided through BMCAC 20418.

10 When, as will be described further below, EBMC 20414 is granted access to MEM 10112 to write data into MEM 10112, data stored in BMCDB 20426 and corresponding addresses from BMCATM 20432 are transferred onto DMID Bus 20430 to DMOVR 20410. As will be described below, DMOVR 20410 then writes that data into those MEM 10112 physical address locations. When data is to be read from MEM 10112 to ED 10124, data is provided by DMOVR 20410 on DMOD Bus 20428 and is transferred into BMCDB 20426.

BMCAC 20418 then reads that data from BMCDB 20426 and transfers that data onto BMCAD Bus 20420 to ED 10124.

During transfers of data from MEM 10112 to ED 10124, MEM .10112 does not provide addresses to be translated into ED 10124 addresses to accompany that data. Instead, those addresses are generated and provided by BMCAC 20418.

NDC 20416 operates in a manner similar to that of EBMC 20414 except that data inputs and outputs of NDCAC 20422 are not buffered through a BMCDB 20426.

As previously described, MEM 10112 has capacity to perform block transfers, that is sequential transfers of .444 44 14

I

4 44 -4- 4 4 -458four 32 bit words at a time. In general, such transfers are performed through EBMC 20414 and are buffered through BMCDB 20426. That is, BMCDB 20426 allows single 32 bit words to be received from ED 10124 by EBMC 20414 and stored therein until a four word block has been received. That block may then be transferred to MEM 10112. Similarly, a block may be received from MEM j 10112, stored in BMCDB 20426, and transferred one word at a time to ED 10124. In contrast, NDC 20416 may generally be utilized for single word transfers.

As indicated in Fig. 204, EBMC 20414, NDC 20416, and each data channel device of IOS 10116 each contain an individual address translation map, for example BMCATM 20432 in EBMC 20414 and NDCATM 20436 in NDC 20416.

Address translation maps stored therein are effectively constructed and controlled by IOCP 20412 for each data channel device. IOS 10116 may thereby provide an individual and separate address translation map for each IOS 10116 data channel device. This allows IOS 10116 to insure that no two data channel devices, nor two groups of data sinks and sources in ED 10124, will mutually einterfere by writing into and destroying data in a common area of MEM 10112 physical address space. Alternately, IOS 10116 may generate address translation maps for two or more data channel devices wherein those maps share a common, or overlapping, area of MEM 10112's physical address space. This allows data stored in MEM 10112 to be transferred between IOS l30116 data channel devices through MEM 10112, and thus to be transferred between L 1 -459various data sink and source devices of ED 10124. For example, a first ED 10124 data source and a first IOS 10116 data channel may write data to be operated upon into a particular area of MEM 10112 address space. The results of CS 10110 operations upon that data may then be written into a common area shared by that first data device and a second data device and read out of MEM 10112 Sto a second ED 10124 data sink by that second data Schannel device. Individual mapping of IOS 10116's data channel devices thereby provides total flexibility in partitioning or sharing of MEM 10112's address space through IOS 10116.

2. -IO- COontf o-l-P-ocesso-2 4--2 As described above, IOCP 20412 is a general purpose computer whose primary function is overall direction and control or data transfer between MEM 10112 and ED 10124.

IOCP 20412 controls mapping of addresses between IOS 10116's data channel devices and MEM 10112 address space. In this regard, IOCP 20412 generates address translation maps for IOS 10116's data channel devices, :*Pe such EBMC 20414 and NDC 20416. IOCP 20412 loads these address translation maps into and controls, for example, BMCATM 20432 of EBMC 20414 and NDCATM 20436 and NDC 20416 through IOCPC Bus 20434. IOCP 20412 also provides certain control functions to DMOVR 20410, as indicated in Fig. 204. In addition to these functions, IOCP 20412 is also provided with data and addressing inputs and outputs. These data addressing inputs and outputs may be utilized, for example, to obtain information utilized by -460- IOCP 20412 in generating and controlling mapping of addresses between IOS 10116's data channel devices and MEM 10112. Also, these data and address inputs and outputs allow IOCP 20412 to operate, in part, as a data channel device. As previously described, IOCP 20412 has data and address inputs and outputs connected from and to DMID Bus 20430 and DMOD Bus 20428. IOCP 20412 thus has access to data being transferred between ED 10124 and MEM 10112, providing IOCP 20412 with direct access to MEM 10 10112 address space. In addition, IOCP 20412 is provided with control and address outputs to NDCAD Bus 20424, thus allowing IOCP 20412 partial control of certain data source and sink devices in ED 10124.

3. Data-Move r-&4M-Ptqf--269 a.a. nput--Dat -Buff-fer- e44&-gn^d SAs described above, DMOVR 20410 comprises an interface between IOS 10116's data channels and MEM 10112. DMOVR 20410 performs actual transfer of data between IOS 10116's data channel devices and MEM 10112, and controls access between IOS 10116's data channel ,devices and MEM 10112. IDB 20440 and ODB 20442 are data and address buffers allowing asynchronous transfer of data between IOS 10116 and MEM 10112. That is, QDB 20442 may accept data from MEM 10112 as that data becomes available and then hold that data until an IOS 10116 data channel device, for example EBMC 20414, is ready to accept that data. IDB 20440 accepts data and MEM r.12 physical addresses from IOS 10116's data channel L -461devices. IDB 20440 holds that data and addresses for.

subsequent transmission to MEM 10112 when MEM 10112 is ready to accept data and addresses. IDB 20440 may, for example, accept a burst, or sequence, of data from EBMC 20414 or single data words from NDC 20416 and' subsequently provide that data to MEM 10112 in block, or four word, transfers as previously described. Similarly, ODB 20442 may accept one or more block transfers or data from ODB 20442 and subsequently provide that data to NDC 20416 as single words, or to DMID 20430 as a data burst.

In addition, as previously described, a block transfer from MEM 10112 may not appear as four sequential words.

In such cases, ODB 20442 accepts the four words of a block transfer as they appear on MIO Bus 10129 and assembles those words into a block comprising four sequential words for subsequent transfer to ED 10124.

Transfer of data through IDB 20440 and ODB 20442 is controlled by PRC 20444, which exchanges control signals with IOCP 20412 and has an interface, previously described, to MEM 10112 through IOMC Bus 10131.

I

bb. rioritvReso-lu r-art- As previously described, PRC 20444 controls access between IOS 10116 data channel devices and MEM 10112.

This operation is performed by means of a Ring Grant Access Generator (RGAG) within PRC 20444.

Referring to Fig. 270e a diagramic representation of PRC 20444's RGAG is shown. In general, PRC 20444's RGAG is comprised of a Ring Grant Code Generator (RGCG) 26910 -462and one or more data channel request comparators. In Fig. 269, PRC 20444's RGAC is shown as including ECLIPSE* Burst Multiplexer Channel Request Comparator (EBMCRC)26912, NOVA* Data Channel Request Comparator (NDCRC) 26914, Data Channel Device X Request Comparator (DCDXRC) 26916, and Data Channel Device Z Request SComparator (DCDZRC) 26918. PRC 20444's RGAG may include r more or fewer request comparators as required by the S« number of data channel devices within a particular IOS S 10 10116.

As indicated in Fig. 269, Request Grant Code (RGC) outputs of RGCG 26910 are connected in parallel to first inputs of EBMCRC 26912, NDCRC 26914, DCDXRC 26916, and DCDZRC 26918. Second inputs of EBMCRC 26912, NDCRC 26914, DCDXRC 26916, and DCDZRC 26918 are connected from other portions of PRC 20444 and receive indications that, respectively, EBMC 20414, NDC 20416, DCDX, or DCDZ has submitted a request for a read or write access to MEM 10112.

Request Grant Outputs (GRANT) of EBMCRC 26912, NDCRC S26914, DCDXRC 26916, and DCDZRC 26918 are in turn I connected to other portions of PRC 20444 circuitry to indicate when read or write access to MEM 10112 has been granted in response to a request by a particular IOS 10116 data channel device. When indication of such a grant is provided to those other portions of PRC 20444, PRC 20444 proceeds to generate appropriate control signals to MEM 10112, through IOMC Bus 10131 as previously described, to IDB 20440 and ODB 20442, and to ;i ".4n -463- IOCP 20412. PRC 20444's control signals initiate that read or write request to that IOS 10116 data channel device. Grant outputs of EBMCRC 26912, NDCRC 26914, DCDXRC 26916, and DCDZRC 26918 are also provided as inputs to RGCG 26910 to indicate, as described further below, when a particular IOS 10116 has requested and been granted access to MEM 10112.

As indicated in Fig. 269, a diagramic figure above RGCG 26910, RGCG generates a repeated sequence of unique 10 RGCs. Herein indicated as numeric digits 0 to 15. Each RGC identifies, or defines, a particular time slot during tr which a IOS 10116 data channel device may be granted access to MEM 10112. Certain RGCs are, effectively, r ^assigned to particular IOS 10116 data channel devices.

ISC 15 Each such data channel device may request access to MEM 10112 during its assigned RGC identified access slots.

St For example, EBMC 20414 is shown as being allowed access t to MEM 10112 during those access slots identified by RGCs t 0, 2, 4, 6, 8, 10, 12, and 14. NDC 20416 is indicated as S 20 being allowed access to MEM 10112 during RGC slots 3, 7, I11, and 15. DCDX is allowed access during slots 1 and 9, and DCDZ is allowed access during RGC slots 5 and 13.

As described above, RGCG generates RGCs 0 to 15 i, a repetitive sequence. During occurrence of a particular S 25 RGC, each request comparator of PRC 20444's RGAG examines that RGC to determine whether its associated data channel device is allowed access during that RGC slot, and whether that associated data channel device has requested access to MEM 10112. If that associated data channel La -464device is allowed access during that RGC slot, and has requested access, that data channel device is granted access as indicated by that request comparator's GRANT output. The request comparators GRANT output is also provided as an input to RGCG 26910 to indicate to RGCG 26910 that access has been granted during that RGC slot.

If a particular data channel device has not claimed and has not been granted access to MEM 10112 during that RGC slot, RGCG 26910 will go directly to next .RGC slot.

10 In next RGC slot, PRC 20444's RGAG again determines n whether the particular data channel device allowed access during that slot has submitted a request, and will grant access if such a request has been made. If not, RGCG 26910 will again proceed directly to next RGC slot, and so on. In this manner, PRC 20444's RGAG insures that each data channel device of IOS 10116 is allowed access to MEM 10112 without undue delay. In addition, PRC 20444's RGAG prevents a single, or more than one, data channel device from monopolizing access to MEM 10112. As S 20 described above, each data channel device is allowed access to MEM 10112 at least once during a particular sequence of RGCs. At the same time, by not pausing within a particular RGC in which no request for access to MEM 10112 has occurred, PRC 20444's RGAG effectively automatically skips over those data channel devices which have not requested access to MEM 10112. PRC 20444's RGAG thereby effectively provides, within a given time interval, more frequent access to those data channel devices which are most busy. In addition, the RGCs -465assigned to particular IOS 10116 data channel devices may be reassigned as required to adapt a particular CS 10110 to the data input and output requirements of a particular CS 10110 configuration. That is, if EBMC 20414 is shown to require less access to MEM 10112 then NDC 20416, certain RGCs may be reassigned from EBMC 20414 to NDC 20416. Access to MEM 10112 by IOS 10116's data channel devices may thereby be optimized as required.

Having described structure and operation of IOS 10 10116, structure and operation of DP 10118 will be described next below.

J t Diagnostc-Ptocesso -Tk-jj--3 3-1 Referring to Fig. 101, as previously described, DP 10118 is interconnected with IOS 10116, MEM 10112, FU 10120, and EU 10122 through DP Bus 10138. DP 10118 is also interconnected, through DPIO Bus 10136, with the external world and in particular with DU 10134. In addition to performing diagnostic and fault monitoring and correction operations, DP 10118 operates, in part, to provide control and display functions allowing an operator to interface with CS 10110. DU 10134 may be comprised, for example, of a CRT and keyboard unit, or a S:r teletype, and provides operators of CS 10110 with all i control and display functions which are conventionally provided by a hard console, that is a console containing switches and lights. For example, DU 10134, through DP 10118, allows an operator to exercise control of CS 10110 for such purposes as system initialization and startup, execution of diagnostic processes, fault monitoring and identification, and control of execution of programs. As *-4 -466will be described further below, these functions are accomplished through DP 10118's interfaces with IOS 10116, MEM 10112, FU 10120, and EU 10122.

DP. 10118 is a general purpose computer system, for example a NOVA* 4 computer of Data General Corporation of Westboro, Massachusetts. Interface of DP 10118 and DU 10134, and mutual operation of DP 10118 and DU 10134, will be readily apparent to one of ordinary skill in the art. DP 10118's interface and operation, with IOS 10116, 10 MEM 10112, FU 10120, and EU 10122 will be described I further next below.

DP 10118, operating as a general purpose computer programmed specificially to perform the functions described above, has, as will be described below, read and write access to registers of IOS 10116, MEM 10112, FU 10120 and'EU 10122 through DP Bus 10138. DP 10118 may read data directly from and write data directly into those registers. As will be described below, these registers are data and instruction registers and are 20 integral parts of CS 10110's circuitry during normal operation of CS 10110. Access to these registers thereby allows DP 10118 to directly control or effect operation 'P of CS 10110. In addition, and as also will be described below, DP 10118 provides, in general, all clock signals to all portions of CS 10110 circuitry and may control operation of that circuitry through control of these clock signals.

For purposes of DP 10118 functions, CS 10110 may be regarded as subdivided into groups of functionally related elements, for example DESP 20210 in FU 10120. DP -467- 10118 obtains access to the registers of these groups, and control of clocks therein, through scan chain circuits, as will be described next below. In general, DP 10118 is provided with one or more scan chain circuits for each major functional sub-element of CS 10110.

Referring to Fig. 205, a diagramic representation of DP 10118 and a typical DP 10118 scan chain is shown. As indicated therein, DP 10118 includes a general purpose Central Processor Unit, or computer, (DPCPU) 27010. A 10 first interface of DPCPU 27010 is with DU 10134 through DPIO Bus 10136. DPCPU 27010 and DU 10134 exchange data t, and control signals through DPIO Bus 10136 in the manner to direct operations of DPCPU 27010, and to display the results of those operations through DU 10134.

15 Associated with DPCPU 27010 is Clock Generator (CLKG) 27012. CLKG 27012 generates, in general, all clock a signals used within CS 10110.

Sl DPCPU 27010 and CLKG 27012 are interfaced with the various scan chain circuits of CS 10110 through DP Bus 20 10138. As described above, CS 10110 may include one or more scan chains for each major sub-element of CS 10110.

One such scan chain, for example DESP 20210 Scan ,Chain (DESPSC) 27014 is illustrated in Fig. 205.

S Interface between DPCPU 27010 and CLKG 27012 and, for 25 example, DESPSC 27014 is provided through DP Bus 10138.

As indicated in Fig. 205, DESPSC 27014 includes Scan Chain Clock Gates (SCCG) 270'5N and one or more Scan Chain Registers (SCRs) 27018 to 27024.

II

;I ~i -468- SCCG 27016 receives clock signals from CLKG 27012 and control signals from DPCPU 27010 through DP Bus 10138.

SCCG 27016 in turn provides appropriate clock signals to the various registers and circuits of, for example, DESP 20210. Clock control signals provided by DPCPU 27010 to SCCG 27016 control, or gate, the various clock signals to these registers and circuits of DESP 20210, thereby effectively allowing DPCPU 27010 to control of DESP 20210.

O SCRs 27018 to 27024 are comprised of various registers within DESP 20210. For example, SCRs 27018 to 27024 may include the output buffer registers of AONGRF 20232, OFFGRF 20234, LENGRF 20236, output registers of OFFALU 20242 and LENALU 20252, and registers within 15 OFFMUX 20240 and BIAS 20246. Such registers are indicated in the present description, as previously described, by arrows appended to ends of those registers, O with a first arrow indicating an input and a second an S* output. In normal CS 10110 operations, as previously 20 described, SCRs 27018 to 27024 operate as parallel in, parallel out buffer registers through which data and instructions are transferred. SCRs 27018 to 27024 are also capable of operating as shift registers and, as indicated in Fig. 205, are connected together to comprise a single shift register circuit having an input from DPCPU 27010 and an output to DPCPU 27010. Control inputs to SCRs 27018 to 27024 from DPCPU 27010 control operation of SCRs 27018 to 27024, that is whether these registers

I

/I

-469shall operate as parallel in, parallel out registers, or as shift registers of DESPSC 27014's scan chain. The shift register scan chain comprising SCRs 27018 to 27024 allows DPCPU 27010 to read the contents of SCRs 27018 to 27024 by shifting the content of these registers into DPCPU 27010. Conversely, DPCPU 27010 may write into SCRs 27018 to 27024 by shifting information generated by DPCPU 27010 from DPCPU 27010 and through the shift register scan chain to selected locations within SCRs 27018 to 10 27024.

Scan chain clock generator circuits and scan chain Sregisters of each scan chain circuit within CS 10110 thereby allow DP 10118 to control operation of each major sub-element of CS 10110. For example, to read information from the scan chain registers therein, and to write information into those scan chain registers as required for diagnostic, monitoring, and control functions.

c Having described structure and operation of each S 20 major element of CS 10110, including MEM 10112, FU 10120, EU 10122, IOS 10116, and DP 10118, certain operations of, in particular, FU 10120 will be described further next C *tr below. The following descriptions will further disclose operational features of JP 10114, and in particular FU 10120, by describing in greater detail certain operations therein by further describing microcode control of JP 10114.

i- -470- F. CS Si(Fias.. 270-2-74) a. Introduction The preceding descriptions have presented the hardware structures and operation of FU 10120 and EU 10122. The following description will describe how devices in FU 10120, and certain EU 10122 devices, function together as a microprogrammable computer, henceforth termed the FU micromachine. The FU i 10 micromachine performs two tasks: it interprets SINs, and IJ it responds to certain signals generated by devices in FU 10120, EU 10122, MEM 10112, and IOS 10116. The signals to which the FU micromachine responds are termed Event signals. In terms of structure and operation, the FU micromachine is characterized by the following: Registers and ALUs specialized for the handling of logical descriptors.

Registers organized as stacks for invocations of microroutines (microinstruction sequences).

S 20 Mechanisms allowing microroutine invocations by means of event signals from hardware.

Mechanisms which allow an invoked microroutine S^ to return either to the microinstruction i following the one which resulted in the invocation or to the microinstruction which resulted in the invocation.

Mechanisms which allow the contents of stack registers to be transferred to MEM 10112, thereby creating a virtual mictostack of limitless size.

p-sl -471- Mechanisms which guarantee response to an event signal within a predictable length of time.

The division of the devices comprising the micromachine into two 'groups: those devices which may be used by all microcode and those which may be used only by KOS (Kernel Operating System, previously described) microcode.

These devices and mechanisms allow the FU micromachine to be used in two ways: as a virtual 10 micromachine and as a monitor micromachine. Both kinds 0* of micromachine use the same devices in FU 10120, but perform different functions and have different logical properties. In the following discussion, when the FU micromachine is being used as a virtual micromachine, it 15 is said to be in virtual mode, and when it is being used as a monitor micromachine, it is said to be in monitor S, mode. Both modes are introduced here and explained in a detail later.

SWhen the FU micromachine is being used in virtual 20 mode, it has the following properties: It runs on an essentially infinite micromachine stack belonging to a Process 610.

It can respond to any number of event signals in *.ret the MO cycle (state) of a single microinstruction.

A page fault may occur on the invocation of any microroutine or on return from any microroutine.

k! 1 -472- When the FU micromachine is in virtual mode, any microroutine may not run to completion, i.e., complete its execution in a predictable length of time, or 'complete it at all.

It is executing a Process 610.

The last four properties are consequences of the first: Event signals result in invocations, and since the micromachine stack is infinite, there is no limit to the number of invocations. The infinite micromachine stack is realized by placing micromachine stack frames on Secure Stack 10336 belonging to a Process 610, and the virtual micromachine therefore always runs on a micromachine stack belonging to some Process 610.

Furthermore, if the invocation of a microroutine or a return from a microroutine requires micromachine frames to be transferred from Secure Stack 10336 to the FU micromachine, a page fault may result, and Process 610 Which is executing the microroutine may be removed from JP 10114, thereby making the time required to execute the microroutine unpredictable. Indeed, if Process 610 is stopped or killed, "he execution of the microroutine may never finish. Ae will be seen in descriptions below, the Virtual Processor 612 is the means by which the virtual micromachine gains access to a Process 610's micromachine stack.

When in monitor mode, the FU micromachine has the following properties: i' It has a micromachine stack of fixed size, the stack is always available to the FU .I ce,. -473micromachine, and it is not associated with a Process 610.

It can respond to only a fixed number of events during the MO cycle of a single microinstruction.

In monitor mode, invocation of a microroutine or return from a microroutine will not cause a page fault.

Microroutines executing on the FU micromachine 10 when the micromachine is in monitor mode are guaranteed to run to completion unless they Sthemselves perform an action which causes them to give up JP 10114.

Microroutines executing in monitor mode need not be performing functions for a Process 610.

Again, the remaining properties are consequences of the first: because the monitor micromachine's stack is of fixed size, the number of events to which the monitor micromachine can respond is limited; furthermore, since 4 20 the stack is always directly accessible to the micromachine, microroutine invocations and returns will not cause page faults, and microroutines running in monitor mode will run to completion unless they themselves perform an action which causes them to give up JP 10114. Finally, the monitor micromachine's stCk is not associated with a Process 610's Secure Stack 10336, and therefore, the monitor micromachine can both execute functions for Processes 610 and execute functions (which -474are related to no Process 610, for example,) the binding and removal of Virtual Processors 612 from JP 10114.

The description which follows first gives an overview of the devices which make up the micromachine, continues with descriptions of invocations on the micromachine and micromachine programming, and concludes with detailed discussions of the virtual and monitor modes and an overview of the relationship between the micromachine and SCS 10110 subsystems. The manner in which the 10 micromachine performs specific operations such as SIN parsing, Name resolution, or address translation may be found in previous descriptions of CS 10110 components which the micromachine uses to perform the operations.

S b. Overview-of-Devices--eomprsn-iS--PU Micromachine.--fFig.--2j7GO Fig. 270 presents an overview of the devices comprising the micromachine. Fig. 270 is based on Fig.

201, but has been simplified to improve the clarity of the discussion. Devices and subdivisions of the micromachine which appear in Fig. 201 have the numbers given them in that figure. When a device in Fig. 270 appears in two subdivisions, it is shared by those subdivisions.

Fig. 270 has four main subdivisions. Three of them are from Fig. 201: FUCTL 20214, which contains the devices used to select the next microinstruction to be executed by the micromachine, DESP 20210, which contains stack and global registers and ALUs for descriptor processing; and MEMINT 20212, which contains the devices which translate Names into logical descriptors and 4 I *I I -475logical descriptors into physical descriptors. The fourth subdivision, EU Interface 27007, represents those portions of EU 10122 which may be manipulated by FU 10120 microcode.

Fig. 270 further subdivides FUCTL 20214 and MEMINT 20212. FUCTL 20214 has four subdivisions: I-Stream Reader 27001, which contains the devices used to obtain SINs and parse them into **4 logic SOPs and Names.

o t10 SOP Decoder 27003, which translates SOPs into f. locations in FU microcode (FUSITT 11012), and in some cases EU microcode (EUSITT 20344), which contain the microcode that performs the I 5 ff corresponding SINs.

Microcode Addressing 27013, which determines the location of the next microinstruction to be executed in FUSITT 11012, S- Register Addressing 27011, which contains t! devices which generate addresses for GRF 10354 registers.

MEMINT 20212 also has three subdivisions: Name Translation Unit 27015, which contains devices which accelerate the translation of Names into logical descriptors.

25 Memory Reference Unit 27017, which contains devices which accelerate the translation of logical descriptors into physical descriptors.

Protection Unit 27019, which contains devices which accelerate primitive access checks on memory references made with logical descriptors.

I

,i -476- Fig. 270 also simplifies the bus structure of Fig.

202 by combining LENGTH Bus 20226, OFFSET Bus 20228, and AONR Bus 20230 into a single structure, Descriptor Bus (DB) 27021. In addition, internal bus connections have been reduced to those necessary for explaining the logical operation of the micromachine. The following discussion first describes those devices used by most microcode executing on FU 10120, and then describes devices used to perform special functions, such as Name 10 translation or protection checking.

1. Devices--used--by--Most--Microcode The subdivisions of the micromachine which contain devices used by most microcode are Microcode Addressing S**e 27013, Register Addressing 27011, DESP 20210, and EU Interface 27007. In addition, most microcode uses MOD Bus 10144, JPD Bus 10142, and DB Bus 27021. The t discussion begins with the buses and then describes the other devices in the above order.

a.a. MOD-Bus -0ol44-;--JPD-Bu--O +4-2--and DB Bus--2-70-2-1 MOD Bus 10144 is the only path by which data may be obtained from MEM 10112. Data on MOD Bus 10144 may have as its destination Instruction Stream Reader 27001, DESP 20210, or EU Interface 27007. In the first case, the data on MOD Bus 10144 consists of SINs; in the second, it is data to be processed by FU 10120, and in the third, it is data to be processed by EU 10122. In the present embodiment, data to be processed by FU 10120 is generally data which is destined for internal use in FU 10120, for example in Name Cache 10226. Data to be processed by EU 10122 is generally operands represented by Names in SINS.

I-

I 4 i -477- JPD Bus 10142 has two uses: it is the path by which data returns to MEM 10112 after it has been processed by JP 10114, and it is the path by which data other than logical descriptors moves between the subdivisions of the micromachine. For example, when CS 10110 is initialized, the microinstructions which are loaded into FUSITT 11012 are transferred from MEM 10112 to DESP 20210 via MOD Bus 101'44, and from DESP 20210 to FUSITT 11012 via JPD Bus 10142.

S 10 DB 27021 is the path by which logical descriptors are transferred in the micromachine. DB 27021 connects Name Translation Unit 27015, DESP 20210, Protection Unit 27019, and Memory Reference Unit 27017. Typically, a ,I Ilogical descriptor is obtained from Name Translation Unit 27015, placed in a register in DESP 20210, and then presented to Protection Unit 27019 and Memory Reference Unit 27017 whenever a reference is made using a logical descriptor. However, DB 27021 is also used to transmit cache entries fabricated in DESP 20210 to ATU 10228, Name t, 20 Cache 10226 and Protection Cache 10234.

b.b. Microcode-Addressing i As discussed here, microcode addressing is comprised of the following devices: Timers 20296, Event Logic "I 20284, RCWS 10358, BRCASE 20278, mPC 20276, MCWO 20292, MCW1 20290, SITTNAS 20286, and FUSITT 11012. All of these devices have already been described in detail, and they are discussed here only as they affect microcode addressing. Other devices contained in Fig. 202, State -478- Registers 20294, Repeat Counter 20280, and PNREG 20282 are not directly relevant to microcode addressing, and are not discussed here.

As has already been described in detail, devices in Microcode Addressing 27013 are loaded from JPD Bus 10142. The microcode addresses provided by these devices and by FUSDT 11010 are transmitted among the devices and to FUSITT 11012 by CSADR Bus 20204. There are six ways o. in which the next microcode address may be obtained: 10 Most commonly, the value in mPC 20276 is incremented, by 1 by a special ALU in mPC 20276, thus yielding the address of the ,microinstruction following the current ,I "microinstruction.

If a microinstruction specifies a call to a microroutine or a branch, the microinstruction contains a literal which an ALU in BRCASE 20278 adds to the value in mPC 20276 to obtain the location of the next microinstruction.

S, 20 If a microinstruction specifies the use of a case value to calculate the location of the next microinstruction, BRCASE 20278 adds a value calculated by DESP 20210 to the value in mPC r 20716. The value calculated by DESP 20210 may be obtained from a field of a logical descriptor, thus allowing the micromachine to branch to different locations in microcode on the basis of type information contained in the L e -479logical descriptor. On return from an invocation of a microroutine, the location at which execution of the microroutine in which the invocation occurred is to continue is obtained from RCWS 10358.

At the beginning of the execution of an SIN, the location at which the microcode for the SIN begins is obtained from the SIN's SOP by means of FUSDT 11010.

SIt, 10 Certain hardware signals cause invocations of Smicroroutines. There are two classes of such l signals: Event signals, which Event Logic 20284 transforms into invocations of certain Smicroroutines, and JAM signals, which are translated directly into locations in microcode.

The addresses obtained as described above are transmitted to SITTNAS 20286, which selects one of the te addresses as the location of the next microinstruction to be executed and transmits the location to FUSITT 11012.

As the location is transmitted to FUSITT 11012, it is also stored in mPC 20276. All addresses except those for Jams are tranferred to SITTNAS 20286 via CSADR Bus 20204. Addresses obtained from JAM signals are transferred by separate lines to SITTNAS 20286.

As will be explained in detail below, microroutine calls and returns also involve pushing and popping micromachine stack frames and saving state contained in MCW1 20290.

-480- Register Addressing 27011 controls access to micromachine registers contained in GRF 10354. As explained in detail below, GRF 10354 contains both registers used for the micromachine stack and global registers, that is, registers that are always accessible to all microroutines. The registers are grouped in frames, and individual registers are addressed by frame number and register number. Register Addressing 27011 allows addressing of any frame and register in the GRs 10360 of GRF 10354, but allows addressing of registers in only three frames of the SR's 10362: the current (top) frame, the previous frame the frame preceding the top frame), and the bottom frame, that is, the lowest frame in a virtual micromachine stack which is still contained in GRF 10354. The values provided by Register Addressing 27011 are stored in MCWO 20292. As will be *explained in the discussion of microroutine invocations S* which follows, current and previous are incremented on S. each invocation and decremented on each return.

c.c Descr r-PCSS -20-2--f a 4.-k DESP 20210 is. a set of devices for storing and processing logical descriptors. The internal structure S of DESP 20210's processing devices has already been explained in detail; here, the discussion deals primarily with the structure and contents of GRF 10354. In a present embodiment of CS 10110, GRF 10354 contains 256 registers. Each register may contain a single logical descriptor. Fig. 271 illustrates a Logical Descriptor -481- 27116 in detail. In a present embodiment of CS 10110, a Logical Descriptor 27116 has four main fields: RS Field 27101, which contains various flags which are explained in detail below.

AON Field 27111, which contains the AON portion of the address of the data item represented by the Logical Descriptor 27116.

OFF Field 27113, which contains the offset portion of the address of the data item represented by Logical Descriptor 27116.

LEN Field 27115, which contains the length of the data item represented by the Logical Descriptor 27116.

'RS Field 27101 has subfields as follows: RTD Field 27103 and WTD Field 27105 may be set by microcode to disable certain Event signals provided for debuggers by CS 10110. For details, see a following description of F debugging aids in CS 10110.

FIU Field 27107 contains two bits. The fields are set from information in the Name Table Entry tused to construct the Logical Descriptor 27116.

The bits determine how the data specified by the Logical Descriptor 27116 is to be justified and filled when it is fetched from MEM 10112.

TYPE Field 27109's four bits are also obtained from the Name Table Entry used to construct the Logical Descriptor 27116. The field's settings vary from S-Language to S-Language, and are used i .it -482to communicate S-Language-specific type information to the S-Language's S-Interpreter microcode.

The four fields of a Logical Descriptor 27116 are contained in three separately-accessible fields in a GRF 10354 register: one containing RS Field 27101 and AON jField 27111, one containing OFF Field 27113, and one containing LEN Field 27115. In addition, each GRF 10354 register may be accessed as a whole. GRF 10354 is 10 further subdivided into 32 frames of eight registers S each. An individual GRF 10354 register is addressed by 1i ,means of its frame number and its register number within the frame. In a present embodiment of CS 10110, half of the frames in GRF 10354 belong to SR's 10362 and are used for micromachine stacks, and half belong to GRs 10360 for storing "global information". In SR's 10362, each GRF 10354 frame contains information belonging to a single invocation of a microroutine. As previously explained, Register Addressing 27011 allows addressing of only three GRF 10354 frames in SR's tack 10362, the current top frame in the stack, the previous frame, and the bottom frame. Registers are accessed by specifying one of these three frames and a register number.

The global information contained in GRs 10360, is information which is not connected with a single invocation. There are three broad categories of global information: Information belonging to Process 610 whose Virtual Processor 612 is currently bound to JP 6 If -483- 10114. Included in this information are the current values of Process 610's ABPs and the pointers which KOS uses to manage Process 610's stacks.

Information required for the operation of KOS.

Included in this information are such items as pointers to KOS data bases which occupy fixed locations in MEM 10112.

Constants, that -is, fixed values required for certain frequently performed operations in FU 10120.

Remaining registers are available to microprogrammers Ias temporary storage areas for data which cannot be stored in a microroutine's stack frame. For example, data which is shared by several microroutines may best be placed in a GR 10360. Addressing of registers in the GRs 10360 of GRF 10354 requires two values: a value of 0 through 15 to specify the frame and a value of 0 through 7 to specify the register in the frame.

As previously discussed in detail, each of the three components AONP 20216, OFFP 20218, and LENP 20220 of DESP S2021.0 also contains ALUs, registers, and logic which allows operations to be performed on individual fields of GRF 10354 registers. In particular, OFFP 20218 contains I 25 OFFALU 20242, which may be used as a general purpose 32 bit arithmetic and logical unit. OFFALU 20242 may further serve as a source and destination for JPD Bus 10142, the offset portion of DB 27021, and NAME Bus 111~111*1 IU C~~L~ -484- 20224, and as a destination for MOD Bus 10144.

Consequently, OFFALU 20242 may be used to perform operations on data on these buses and to transfer data from one bus to another. For example, when an SIN contains a literal value used in address calculation, the literal value is transferred via NAME Bus 20224 to OFFALU 20242, operated on, and output via the offset portion of DB 27021.

EU 10122-2-Interface 10 FU 10120 specifies what operation EU 10122 is to perform, what operands it is to perform it on, and when it is finished, what is to be done with the operands. FU 10120 can use two devices in EU 10122 as destinations for data, and one device as a source for data. The destinations are COMQ 20342 and OPB 20322. COMQ 20342 receives the location in EUSITT 20344 of the microcode which is to perform the operation desired by the FU 10120. COMQ 20342 may receive the location in microcode either from an FU 10120 microroutine or from an SIN's 1 20 SOP. In the first case, the location is transferred via JPD Bus 10142, and in the second, it is obtained from EUSDT 20266 and transferred via EUDIS Bus 20206. OPB 20322 receives the operands upon which the operation is S to be performed. If the operands come directly from MEM 10112, they are transferred to OPB 20322 via MOD Bus 10144; if they come from registers or devices in FU 10120, they are transferred via JPD Bus 10142.

Result Register 27013 is a source for data. After EU 10122 has completed an operation, FU 10120 obtains the result from Result Register 27013. FU 10120 may then f1* I. -485place the result in MEM 10112 or in any device accessible from JPD Bus 10142.

2. Specia-i+zed-Miciomach-ine---ev+ces Each of the groups of specialized devices serves one of CS 10110's subsystems. I-Stream Reader 27001 is part of the S-Interpreter subsystem, Name Translation Unit 27015 is part of the Name Interpeter subsystem, Memory Reference Unit 27017 is part of the Virtual Memory ##00 $Management System, and Protection Unit 27019 is part of 10 the Access Control System. Here, these devices are 4 so S. explained only in the context of the micromachine; for a complete understanding of their functions within the subsystems to which they belong, see previous *descriptions of the subsystems.

a.a. *I-Stream-Reader--2700.1 I-Stream Reader 27001 reads and parses a stream of t SINs (termed the I-Stream) from a Procedure Object 604, 606, 608. The I-Stream consists of SOPs (operation codes), Names, and literals. As previously mentioned, in a present embodiment of CS 10110, the i-Str-eam read from a given Procedure 602 has a fixed format: the SOPs are 8 S bits long and the Names and literals all have a single length. Depending on the procedure, the length may be 8, S12, or 16 bits. I-Stream Reader 27001 parses the I- Stream by breaking it up into its constituent SOPs and Names and passing the SOPs and Names to appropriate parts of the micromachine. I-stream Reader 27001 qcntain two groups of devices: PC Values 27006, which is made up of trb-' registers which contain locations in S.4 r -486- Stream. When added to ABP PBP, the values contained in these registers specify locations in Procedure Object 901 containing the Procedure 602 being executed. CPC 20270 contains the location of the SOP or Name currently being interpreted; IPC 20272 contains the location of the beginning of the SIN currently being executed; EPC 20274, finally, is of interest Ireit' only at the beginning of the execution of an 10 SIN; at that time, it contains the location of Sthe last SIN to be executed.

Parsing Unit 27005, which is made up of INSTB S20262, PARSER 20264, and PREF 20260. The Smicromachine uses PREF 20260 to create Logical Descriptors 27116 for the I-Stream, which are then placed on DB Bus 27021 and used in logical memory references. The data returned from these references is placed in INSTB 20262, and parsed Sby PARSER 20264.

SOPs, Names, and literals obtained by PARSER 20264 are placed on NAME Bus 20224, which connects PARSER I20264, SOP Decoder 27003, Name Translation Unit 27015, and OFFALU 20242.

b.b. SOP-Decoder--2700 25 SOP Decoder 27003 decodes SOPs into locations in FP 10120 and EU 10122 microcode. SOP Decoder 27003 vt comprises FUSDT 11010, EUSDT 20266, Dialect Register (RDIAL) 24212, and LOPDCODE 24210. FUSDT 11010 are icr mac in us s P EF 026 t cr ate Log ca 1 Desc.2..~ ~yrpto iICr.- 2711 fo the Stram whc are~I -487further comprised of FUDISP 24218 and FALG 24220. The manner in which these devices translate SOPs contained in SINs into locations in FUSITT 11012 and EUSITT 20344 has been previously described.

c.c. Name- Trans-lation-Un-i-27--15 Name Translation Unit 27015 accelerates the translation of Names into Logical Descriptors 27116.

This operation is termed name resolution. It is comprised of two components: NC 10226 and Name Trap 10 20254. NC 10226 contains copies of information from a t Procedure Object 604's Name Table 10350, and thereby makes it possible to translate Names into Logical Descriptors 27116 without referring to Name Table 10350.

When a Name is presented to Name Translation Unit 27015, it is latched into Name Trap 20254 for later use by Name Translation Unit 27015 if required. As will be explained in detail later, in the present embodiment, Name translation always begins with the presentation of a Name to NC 10226. If the Name has already been translated, the information required to construct its Logical Descriptor 27116 may be contained in NC 10226. If there is no information for the Name in NC 10226, Name i Resolution Microcode obtains the Name from Name Trap 20254, uses information from Name Table 10350 for the procedure being executed to translate the Name, places the required information in NC 10226, and attempts the translation again. When the translation succeeds, a Logical Descriptor 27116 corresponding to the Name is produced from the information in Name Cache 10115, placed on DB Bus 27021, and loaded into a GRF 10354 register.

-488d.d. Memorv Reference--Unit--27Q0IM Memory Reference Unit 27017 performs memory references using Logical Descriptors 27116. Memory l Reference Unit 27017 receives a command for MEM 10112 and a Logical Descriptor 27116 describing the data upon which the command is to be performed. In the case of a write operation, Memory Reference Unit 27017 also receives the data being written via JPD Bus 10142. Memory Reference Unit 27017 translates Logical Descriptor 27116 to a 10 physical descriptor and transfers the physical descriptor and the command to MEM 10112 via PD Bus 10146. A Memory Reference Unit 27017 has four components: ATU 10228, which contains copies of information from KOS virtual r t f memory management system tables, and thereby accelerates logical-to-physical descriptor translation; Descriptor Trap 20256, which traps Logical Descriptors 27116, Command Trap 27018, which traps memory commands; and Data Trap 20258, which traps data on write operations. When a logical memory reference is made, a Logical Descriptor 27116 is presented via DB Bus 27021 to ATU 10228, and at the same time, Logical Descriptor 27116 and the memory command are trapped in Descriptor Trap 20256 and Command S Trap 27018. On write operations, the data to be written is trapped in Data Trap 20258. If the information needed to form the physical descriptor is present in ATU 10228, the physical descriptor is transferred to MEM 1C112 via PD Bus 10146. If the information needed to form the physical descriptor is not present in ATU 10228, an Event Signal from ATU 10228 invokes a microroutine which -489retrieves Logical Descriptor 27116 from Descriptor Trap 20256 and uses information contained in KOS virtual memory management system tables to make an entry in ATU 10228 for Logical Descriptor 27116. When the microroutine returns, the logical memory reference is repeated using Logical Descriptor 27116 from Descriptor Trap 20256, the memory command from Command Trap 27018, and on write operations, the data in Data Trap 20258. As will be described in detail in the discussion of virtual memory management, if the data referenced by a logical memory reference is not present in MEM 10112, the logical memory reference causes a page fault.

e e.e. The-Protection-Unf--27U-o9 On each logical memory reference, Protection Unit 27019 checks whether the subject making the reference has -access rights which allow it to perform the action 't specified by the memory command on the object being referenced. If the subject does not have the required access rights, a signal from Protection Unit 27019 causes MEM 10112 to abort the logical memory reference.

Protection Unit 27019 consists of Protection Cache 10234, which contains copies of information from KOS Access Control System tables, and thereby speeds up protection checking, and shares Descriptor Trap 20256, Command Trap 27018, and Data Trap 20258 with Memory Reference Unit 27017. When a logical memory reference is made, the AON and offset portions of the logical descriptor are presented to Protection Cache 10234. If Protection Cache -490- 10234 contains protection information for the object specified by the AON and offset and the subject performing the memory reference has the required access, the memory reference may continue; if Protection Cache 10234 contains protection information and the subject does not have the required access, a signal from Protection Cache 10234 aborts the memory reference. If Protection Cache 10234 does not contain the required "r \access information, a signal from Protection Cache 10234 10 aborts the memory reference and invokes a microroutine S..which obtains the access information from KOS Access i Control System tables and places it in Protection Cache S10234. When Protection Cache 10234 is ready, the memory tr access is repeated, using the logical descriptor from Descriptor Trap 20256, the memory command from Command Trap 27018, and in the case of write operations, the data I in Data Trap 20258.

f.f. KOS -Micromachin--Devces iAs mentioned in the above introduction to the micromachine, the devices making up the micromachine may be divided into two classes: those which any microcode .j written for the micromachine may manipulate, and those which may be manipulated exclusively by KOS microcode.

The latter class consists of certain registers in GRs 10360 of GRF 10354, the bottom frame of the portion of the virtual micromachine stack in the stack portion (Stack Registers 10362) of GRF 10354, and the devices contained in Protection Unit 27019 and Memory Reference Unit 27017. Because Protection Unit 27019 and Memory _o -491- Reference Unit 27017 may be manipulated only by KOS microcode, non-KOS microcode may not use Descriptor Trap 20256 or Command Trap 27018 as a source or destination, may not load or invalidate registers in ATU 10228 or Protection Cache 10234, and may not perform physical memory references, memory references which place physical descriptors directly on PD Bus 10146, instead of presenting logical descriptors to Memory Reference Unit 27017 and Protection Unit 27019. Similarly, non-KOS 10 microcode may not specify KOS registers in the GRs 10360 of GRF 10354 or the bottom frame of the stack portion of GRF 10354 when addressing GRF 10354 registers. Further, S' in embodiments allowing dynamic loading of FUSITT 11012, only KOS microcode may manipulate the devices provided for dynamic loading.

In a present embodiment of CS 10110, the distinction between KOS devices and registers and devices and registers accessible to all microprograms is maintained by the microbinder. The microbinder checks all microcode for microinstructions which manipulate devices in Protection Unit 27019, or Memory Reference Unit 27017, or which address GRF 10354 registers reserved for KOS use.

However, it is characteristic of the micromachine that KOS devices are logically and physically separate from devices accessible to all microprograms and, consequently, other embodiments of CS 10110 may use hardware devices to prevent non-KOS microprograms from manipulating KOS devices.

-492c. M-i croma chJine-Stacs--an--M iT+t e-C-a-3-s and- Ret urns Mi r omachine-S tacks--+fP+--o--2- As previously mentioned, the FU micromachine is a stack micromachine. The properties of the FU micromachine's stack depends on whether the FU micromachine is in virtual or monitor mode. In virtual mode, the micromachine stack is of essentially unlimited S. size; if it contains more frames than allowed for inside S 10 FU 10120, the top frames are in GRF 10354 and the remaining frames are in Secure Stack 10336 belonging to Process 610 being executed by the FU micromachine. In the following, the virtual mode micromachine stack is C:fV termed the virtual micromachine stack. In monitor mode, the micromachine stack consists of a fixed amount of storage; in a present embodiment of CS 10110, the monitor mode micromachine stack is completely contained in the t stack portion, SRs 10362, of GRF 10354; in other 4 tembodiments of CS 10110, part or all of the monitor mode micromachine stack may be contained in an area of MEM 10112 which has a fixed size and a fixed location known /to the monitor micromachine. In yet other embodiments of CS 10110, monitor mode micromachine stack may be of flexible depth in a manner similar to the virtual micromachine stack. In either mode, microroutines other t than certain KOS microroutines which execute state save and restore operations may access only two frames of GRF 10354 stack: the frame upon which the microroutine is executing, called the current frame, and the frame upon -493which the microroutine that invoked that microroutine executed, called the previous frame. KOS microroutines which execute state save and restore operations may in addition access the bottom frame of that portion of the virtual micromachine stack which is contained in GRF 10354.

Fig. 272 illustrates stacks for the FU micromachine.

Those portions of the micromachine stack which are contained in the FU are contained in SR's 10362 (of GRF 10354) and in RCWS 10358. Each register of RCWS 10358 is t permanently associated with a GRF frame in SRs 10362 of J GRF 10354, and the RCWS 10358 register and the GRF frame together may contain one frame of a micromachine stack.

As previously describe, each register of GRF 10354 contains three fields: one for an AON and other information, one for an offset, and one for a length. As illustrated in Fig. 251, each register in RCWS 10358 contains four fields: A one bit field which retains the value of the Condition Code register in MCW1 20290 at the time that the invocation which created the next 4r* *frame occurred.

A field indicating what Event Signals were Spending at the time that the invocation to which the RCWS register belongs invoked another microroutine.

A flag indicating whether the microinstruction being executed when the invocation occurred was the first microinstruction in an SIN.

4 4' -494- The address at which the execution of the invoking microroutine is to continue.

The uses of these fields will become apparent in the ensuing discussion.

The space available for micromachine stacks in SRs 10362 and RCWS 10358 is divided into two parts: Frames 27205 reserved for MOS 10370 and Frames 27206 available for the MIS 27203. Frames 27206 may contain no MIS e t* Frames 27203, or be partially or completely occupied by t 1 10 MIS Frames 27203. Space which contains no MIS Frames is SFree Frames 27207. The size of the space reserved for Monitor Micromachine Stack Frames 27205 is fixed, and Spaces 27203, 27205, and 27207 always come in the specified order. Register Addressing 27011 handles addressing in Stack Portion 27201 of GRF 10354 and RCWS 10358 in such fashion that the values for the locations of current, previous, and bottom frames specifying registers in RCWS 10358 or frames in Stack Portion 27201 t, automatically "wrap around" when they are incremented beyond the largest index value allowed by the sizes of the registers or decremented below the smallest index H t value. Thus, though Spaces 27203, 27205, and 27207 always have the same relative order, their GRF 10354 frames and RCWS registers may be located anywhere in Stack Portion 27201 and RCWS 10358.

2. Micro outine--nvoca±ons-and-Reterrs In CS 10110, microroutines may be invoked by other microroutineS or by signals from CS 10110 hardware. The methods of invocation aside, microroutine invocations and -495returns resemble invocations of and returns from procedures written in high-level languages. In the following, the general principles of microroutine invocations and returns are discussed, and thereafter, the specific methods by which microroutines may be invoked in CS 10110. The differences between invocations in monitor mode and invocations in virtual mode are explained in the detailed discussions of the two modes.

The microroutine which is currently being executed 10 runs on the frame specified by Current Pointer 27215.

When an invocation occurs, either because the executing ,,microroutine performs a call, or because a signal which t causes invocations has occured, JP 10114 hardware does f *p three things: It stores state information for the invoking Smicroroutine in the RCWS 10358 register associated with the current frame. The state information includes the location at which execution of the invoking microroutine will resume, as well as other state information.

It increments Current Pointer 27215 and Previous T Pointer 27213, thereby providing a frame for the new invocation.

It begins executing the first instruction of the eowly invoked microroutine.

Because the newly-invoked microroutine can access registers of the invoking microroutine's frame, the invoking microroutine can pass "arguments" to the invoked microroutine by placing values in registers in its frame I

I-

h. -496used by the invoked microroutine. However, the invoking microroutine cannot specify which registers contain "arguments" on an invocation, so the invoked microroutine must know which registers of the previous frame are used by the invoking microroutine. Since the only "arguments" which a micrrooutine has access to are those in the previous frame, a microroutine can pass arguments which it received from its invoker to a microroutine which it invokes only by copying the arguments from its invoker's 10 frame to its own frame, which then becomes the newlyinvoked routine's previous frame.

The return is the reverse of the above: Current Pointer 27215 and Previous Pointer 27213 are decremented, ,E thereby "popping off" the finished invocation's frame and returning to the invoker's frame. The invoker then resumes execution at the location specified in the RCWS 10358 register and using the state saved in the RCWS 1 10358. The saved state includes the value of the S, Condition Code in MCW1 20290 at the time of the S 20 invocation and flags indicating various pending Events.

The Condition Code field in MCW1 20290 is set to the saved value, and the pending event flags may cause Events to occur as described in detail below.

3. Means- of -Invoking-Micrfoti*tes In the micromachine, invocations may be produced either by commands in microinstructions or by hardware signals. In the following, invocations produced by commands in microinstructions are termed Calls, while those produced by hardware signals are termed Event invocations and Jams. Invocations are further distinguished from each other by the locations to which they return. Calls and Jams return to the microinstruction following the microinstruction in which the invocation occurs; Event invocations return to that microinstruction, which is then repeated.

In terms of implementation, the different return -t -Ilocations are a consequence of the point in the 44 10 micromachine cycle at which Calls, Jams, and Event invocations save a return location and transfer control to the called routine. With Calls and Jams, these Soperations are performed in the Ml cycle; with Event invocations, on the other hand, the Event signal during the MO cycle causes the MO cycle to be followed by a MA cycle instead of the M cycle, and the operations are performed in the MA cycle. In the M1 cycle, the value in mPC 20276 is incremented; in the MA cycle, it is not.

t I Consequently, the return value saved in RCWS 10358 on a Call or Jam is the incremented value of mPC 20276, while th return value saved on an Event invocation is the aunincremented value of mPC 20276. The following discussion will deal first with Calls and Jams, and then with Event invocations.

A Call command in a mictoinstruction contains a literal value which specifies the offset from the microinstruction containing the Call at which execution is to continue after the Call. When the microinstruction with the Call command is executed in micromachine cycle M,1 BRCASE 20278 adds the offset contained in the command fec eunvlesvdo a vn noaini h -498to the current value of mPC 20276 in order to obtain the location of the invoked microroutine and sets SITTNAS 20286 to select the location provided by BRCASE 20278 as the location of the next microinstruction. Then the Call command increments mPC 20276 and stores the incremented value of mPC 20276 in the RCWS 10358 register associated with the current frame in SRs 10362 and increments Current Pointer 27215 and Previous Pointer 27213 to provide a new frame in SRs 10362. The Jam works exactly 10 like the Call, except that a hardware signal during t ,It micromachine cycle Ml causes the actions associated with the invocation to occur and provides the location of the y invoked microroutine directly to SITTNAS 20286.

With Events, Event Logic 20284 causes an invocation to occur during cycle MO and provides the location of the invoked microroutine via CSADR 20299. Since the Event occurs during cycle MO, the location stored in RCWS 10358 is the unincremented value of mPC 20276, and SITTNAS 20286 selects the location provided by Event Logic 20284 as the location of the next microinstruction. Since the return from the Event causes the microinstruction during 1 which the Event occurred to be re-executed, the t microinstruction and the microroutine to which it belongs may be said to be "unaware" of the Event's occurrence.

The only difference between the execution of a Smicrinstruction during which an Event occurs and the execution of the same microinstruction without the Event is the length of time required for the execution.

r -499- 4. Occur renc e--of-Bvent-nvocations--+tP As described previously, Event invocations are produced by Event Logic 20284. The location in microcode to which Event Logic 20284 transfers control is determined by the following: The operation being commenced by FU 10120.

Certain Event invocations may occur only at the t.r. beginning of 'certain FU 10120 operations.

The state of Event signal lines from hardware ,and internal registers in Event Logic 20284.

The state of certain registers visible via MCW1 S20290. Some of these registers enable Events and others mask Events. Of the registers which enable Events, some are set by Event signals and others by the microprogram.

On returns from invocations of microroutines, the settings of certain bits in the RCWS 10358 I register belonging to the micromachine frame for the invocation that is being returned to.

Microprograms may use these mechanisms to disable Event signals and to delay an Event invocation from an t ttt Event signal for a single microinstruction or an indefinite period, and FU 10120 uses them to 25 automatically delay Event invocations resulting from certain Event signals. Using traditional programming terminology, the mechanisms allow a differential masking of Event signals. An Event 4ignal may be explicitly masked for a single microinstruction, it may be masked hb -500for a sequence of microinstructions, it may be automatically masked until a certain operation occurs, or it may be automatically masked for a certain maximum length of time. Event signals which occur while they are masked are not lost. In some cases, the Event signal continues until it is serviced; in others, a register is set to retain the fact that the Event signal occurred.

When the Event signal is unmasked, the set register causes the Event signal to reoccur. In some cases, finally, the Event signal is not retained, but recurs when the microinstruction which caused it is repeated.

In the following, the relationship between FU 10120 operations and Event signals is first presented, and then a detailed discussion of the enabling registers in MCW1 20290 and of the bits in RCWS 10358 registers which control Event invocations.

FU 10120 allows Event invocations resulting from Is .Event signals to be inhibited for a single i, microinstruction; it also delays certain Event 20 invocations for certain Event signals until the first microinstruction of an SIN. Other Event signals occur ~only at the beginning of an SIN, at the beginning of a Namespace Resolve or Evaluate operation, or at the Sbeginning of a logical memory reference.

Event invocations may be delayed for a single microinstruction by setting a field of the microinstruction itself. Setting this field delays almost all Evcnt ,nvocations, and thereby guarantees that an Event invocation will not occur during the microinstruction's MO cycle.

Whn te Evnt igna isunmsked th setregste cass'eEvn inltorocr n oecss s-l l t I 'I~ -501- Event signals relating to debugging occur at the beginnings of certain micromachine operations. Such Event signals are called Trace Event signals. As will be explained in detail, in the discussion of the debugger, Trace Event signals can occur on the first microinstruction of an SIN, at the beginning of an Evaluate or Resolve operation, at the beginning of a logical memory reference, or at the beginning of a microinstruction. IPM interrupt signals and Interval Timer Overflow Event signals are automatically masked Ec" until the beginning of the next SIN or until a maximum l t amount of time has elapsed, which ever occurs first. The mechanisms involved here are explained in detail in the discussion of interrupt handling in the FU 10120 micromachine.

Turning now to the registers used to mask and enable Event signals, Fig. 273 is a representation of the masking and enabling registers in MCW1 20290 and of the field in RCWS 10358 registers which controls Event 20 invocations. Beginning with the registers in MCW1 20290, there are three registers which control Event invocations: Event Mask Register (EM) 27301, Events k Pending Register (EP) 27309, and Trace Enable Register

S

t I. (TE) 27319. Bits in EM 27301 mask certain Event signals as long as they are set; bits in EP Register 27309 record the occurrence of certain Event signals while they are masked; when bits in TE Register 27319 are set, Trace Event signals occur before certain FU 10120 operations.

-502- EM 27301 contains three one bit fields: Asynchronous Mask Field 27303, Monitor Mask Field 27305, and Trace Event Mask Field 27307. As explained in detail in the discussion of FU 10120 hardware, these bits establish a hierarchy of Event masks. If Asynchronous Mask Field 27303 is set, only two Event signals are masked: that resulting from an overflow of EGGTMR 25412 and that resulting from an overflow of EU 10122's stack. If Monitor Mask Field 27305 is set, those Events are masked, S 10 and additionally, the FU Stack Overflow Event signal is masked. As will be explained in detail later, when the t r FU 10120. Stack Overflow Event signal is masked, the FU H micromachine is executing in monitor mode. If Trace Event Mask Field 27307 is set, Trace Trap Event signals are masked in addition to the above signals. Each of the fields in EM 27301 may be individually set and cleared by the microprogram.

Four Event signals set fields in EP 27309: the EGGTMR 25412 Runout signal sets ET Field 27311, the INTTMR 25410 S 20 Runout signal sets IT Field 27313, the Non-Fatal Memory Error signal sets ME Field 27315, and the Inter-Process Message signal sets IPM Field 27317. Event invocations t 'for all of these Event signals but the Egg Timer Runout signal occur at the beginning of an SIN; in these cases the fields in EP 27309 retain the fact that the Event signal has occurred until that time; the Event invocation for the Egg Timer Runout signal occurs as soon after the signal as the settings of mask bits in EM 27301 allow.

The bit in ET Field 27311 retains the fact of the Egg i -503- Timer Runout signal until the masking allows the Event invocation to occur. All of the fields in EP 27309 but ME Field 27315 may be reset by microcode. The microroutines invoked by the Events must reset the appropriate fields; otherwise, they will be reinvoked when they return. ME Field 27315 is automatically reset when the memory error is serviced.

TE Register Field 27319 enables tracing. Each bit in the register enables a kind of Trace Event signal when it S 10 is set. Depending on the kind of tracing, the Trace Event Ssignal occurs at the beginning of an SIN, at the tot t beginning of a Resolve or Evaluate operation, at the beginning of a logical'memory reference, or at the t t*o beginning of a microinstruction. For details, see the following description of debugging.

Turning now to the registers contained in RCWS 10358, each RCWS Register 27322 contains eight fields which control Event signals. The first field is FM Field 27323. FM Field 27323 reflects the value of a register 20 in Event Logic 20284 when the invocation to which RCWS Register 27322 belongs occurs. The register in Event Logic 20284 is set only when the microinstruction currently being executed is the first microinstruction of *an SIN. Thus, FM Field 27323 is set only in RCWS Registers 27322 belonging to Event invocations which occur in the MO cycle of the first microinstruction in the SIN, at the beginning of the SIN. The value of the register in Event Logic 20284 is saved in FM Field 27323 because several Event invocations may occur at the j j i i 1 t -504beginning of a single SIN. The Event invocations occur in order of priority: when the one with the highest priority returns, the fact that FM Field 27323 is set causes the register in Event Logic 20284 to again be set 5 to the state which it has on the first microinstruction of an SIN. The register's state, thus set, causes the next Event invocation which must occur at the beginning of the SIN to take place. After all such invocations are i finished, the first microinstruction enters its Ml cycle 10 and resets the register in Event Logic 20284. In its reset state, the register inhibits all Event invocations which may occur only at the beginning of an SIN. It is again set at the beginning of the next SIN.

Ct aC The remaining fields in RCWS Register 27322 which control Event invocations are the fields in Return Signals Field 27331. These fields allow the information C cc that an Event signal has occurred to be retained through t zEvent invocations until the Event signal's Event invocation takes place. When an invocation occurs, these i 20 fields are set by Event Logic 20284. On return from the invocation, the values of the fields are input into Event Logic 20284, thereby producing Event signals. The Event signal with the highest priority results in an Event invocation, and the remaining Event signals set fields in Return Signals Field 27331 belonging to RCWS Register 27322 belonging to the invocation which is being executed when the Event signals occur. Because the fields in i Return Signals Field 27330 are input into Event Logic I 20284, microcode invoked as a consequence of Event signals which sets one of these fields must reset the I! i 1 i it^s 1 -505field itself. Otherwise, the return from the microcode will simply result in a reinvocation of the microcode.

The seven fields in Return Signals Field 27330 have the following significance: When EG Field 27333 is set, an EU 10122 dispatch operation produced an illegal location in EU 10122 microcode EUSITT 20344.

When NT Field 27335, ST Field 27341, mT Field 27343, or mB Field 27345 is set, a trace signal 10 has occurred. These are explained in detail in the discussion of debugging.

VCf C When ES Field 27337 is set, an EU 10122 c Storeback Exception has occurred, an error C occurred when EU 10122 attempted to store the result of an operation in MEM 10112.

S- When MRR Field 27339 is set, a condition such as CV an ATU 10228 miss or a Protection Cache 10234 miss has occurred, and it is necessary to S: t reattempt a memory reference.

d. Virtual--Micromach-ines-and--thqM-coct-n Micromach-ine As previously described, microcode being executed on II.: 4FU 10120's micromachine can run in either monitor mode or virtual mode. In this portion of the discussion, the distinguishing features and applications of the two modes are explained in detail.

1. Vttia:u -Mode

I

As previously mentioned, the chief distinction between virtual mode and monitor mode is MIS 10368. The Ii t l T i 11 1 l l I -506fact that MIS 10368 is of essentially unlimited size has the following consequences for microroutines which execute in virtual mode.

An invocation of a microroutine executing in virtual mode may have as its consequence further invocations to any depth.

Any invocation of or return from a microroutine executing in virtual mode may cause a page cl fault.

The FU micromachine is in virtual mode when all bits in the Event Masks portion of MCW1 20290 are cleared. In this state, no enabled Event signals are masked, and Event invocations may occur in any microinstruction which does not itself mask them.

Because invocations may occur to any depth in virtual mode, microroutines executing in this mode may be recursive. Such recursive microroutines are especially useful for the interpretation of Names. Often, as previously described, the Name Table Entry for a Name will contain Names which resolve to other Names, and the virtual micromachine's limitless stack allows the use of recursive Name Resolution microroutines in such situations. Recursive microroutines may also be used for •complex SINs, such as Calls.

Because invocations can occur to any depth, any number of Events may occur while a microroutine is executing in monitor mode. This in turn greatly simplifies Event handling. If an Event signal occurs while an Event with a given priority is being handled and the Event being signalled has a higher priority than the -507one being handled, the result is simply the invocation of the new Event's handler. Thus, the order in which the Event handlers finish corresponds exactly to the priorities of their Events: those with the highest finish first.

A page fault may occur on any microinvocation or return executed in virtual mode because an invocation in virtual mode which occurs when there are no more Free Frames 27207 on SRs 10362 causes an Event signal which invokes a microroutine running in monitor mode. The microroutine transfers MIS Frames 27203 from GRF 10354 to SSecure Stack 10336 in MEM 10112, and the transfer may *pi cause a page fault. Similarly, when a microreturn takes 7 place from the last frame on MIS Frames 27203 on SRs 10362, an Event signal occurs which invokes a microroutine that transfers additional frames from Secure Stack 10336 to GRF 10354, and this transfer, too, may cause a page fault.

The fact that page faults may occur on S 20 microinvocations or microreturns in virtual mode has two important consequences: microroutines which cannot tolerate page faults other than those explicitly tgenerated by the microroutine itself cannot execute in S" virtual mode, and because unexpected page faults cause execution to become indeterminate, microroutines which must run to completion cannot execute in virtual mode.

For example, if the microroutine which handles page faults executed in virtual mode, its invocation could cause a page fault, which would cause the microroutine to -A ,k -508be invoked again, which would cause another page fault, and so on through an infinite series of recursions.

2. Monito- -Micromech-ine As previously described, the essential feature of monitor mode is MOS 10370. In a present embodiment of CS 10110, this stack has a fixed minimum size, and is always contained in GRF Registers 10354. The nature of MOS 10370 has four consequences for microroutines which execute in monitor mode: When the micromachine is in monitor mode, the depth of invocations is limited; recursive microroutines therefore cannot be executed in monitor mode, and Event invocations must be I ^S limited.

Invocations of microroutines or returns from microroutines in monitor mode never result in Spage faults.

Microroutines executing in monitor mode are guaranteed to run to completion if they do not suspend the Process 610 which they are executing or perform a Call to software.

S When the micromachine is executing in monitor t mode, it is guaranteed to return to virtual mode t within a reasonable period of time, either because a microroutine executing in monitor mode has run to completion, or because the microroutine has suspended the Process 610 which it is executing, or has made a Call to software. The result in both cases is the 4 l j j -509execution of a new sequence of SOPs, and thus a return to virtual mode.

In a present embodiment of CS 10110, the FU micromachine is in monitor mode when a combination of masking bits in MCW1 20290 is set which results in the masking of the FU Stack Overflow Event and the Egg Timer Overflow Event. As previously described, these Events are masked if Fields 27303, 27305, or 27307 is set.

These Events and the consequences of masking them are 10 explained in detail below.

SThe event signal for the FU Stack Overflow Event occurs on microinvocations for which there is no frame t r" available in MIS Frames 27203. If the Event signal is not masked, it causes, the invocation of a microroutine which moves MIS Frames from MIS Frames 27203 onto a Process 610's Secure Stack 10336. When the FU Stack Overflow Event is masked, all frames in SRs 10362 of GRs S, 10360 are available for microroutine invocations and 1, microroutine invocations will not result in page faults, but if the capacity of SRs 10362 is exceeded, FU 10120 ceases operation.

1 The Egg Timer Overflow event signal occurs when Egg STMR 25412 runs out. As will be explained in detail later, Egg TMR 25412 ensures that an Interval Timer Runout, an Inter-processor Message, or a Non-fatal Memory Error will be serviced by JP 10114 within a reasonable amount of time. If an Interval Timer Runout Event signal or an Inter-processor Message Event signal occurs at a time when it is inefficient for the FU micromachine to handle the Event, Egg TMR 25412 begins running. When Egg L -510- TMR 25412 runs out, the Event is handled unless the micromachine is in monitor mode. If the Egg TMR 25412 Runout Event signal occurs while the FU micromachine is in monitor mode, while the Event is masked, the Event signal sets Field 27311 in MCW1 20290. When the FU micromachine reverts to virtual mode, when all Event Mask bits in MCWI 20290 are cleared, the Egg TMR 25412 Runout Event occurs, and the Interval Timer Runout and/or the Inter-processor Message Event handlers are invoked by Event Logic 20284.

e. -InterfuPt--and--Fau-t-t-Hand-+nqg g 1. General--Princip-1es Any computer system must be able to deal with with occurrences which disrupt the normal execution of a program. Such occurrences are generally divided into two classes: faults and interrupts. A fault occurs as a consequence of an attempt to execute a machine i instruction, and its occurrence is therefore synchronous with the machine instruction. Typical faults are

S

t 20 floating point overflow faults and page faults.

A

floating point overflow fault occurs when a machine instruction attempts to perform a floating point S arithmetic operation and the result exceeds the capacity of the CS 10110's floating point hardware, that is EU 10122. A page fault occurs when a machine instruction in a computer system with virtual memory attempts to reference data which is not presently available in the computer system's primary memory, that is MEM 10112.

Since faults are synchronous with the execution of machine instructions and in many cases the result of the L S f -511execution of specific machine instructions, their occurrence is to some extent predictable.

The occurrence of an interrupt is not predictable.

An interrupt occurs as a consequence of some action taken by the computer system which has no direct connection with the execution of a machine instruction by the computer system. For example, an I/O interrupt occurs when data transmitted by an I/O device (IOS 10116) *reaches the central processing unit (FU 10120), 10 regardless of the machine instruction the central processing unit is currently executing.

In conventional systems, interrupts and faults have been handled as follows: if an interrupt or fault occurs, j 'the computer system recognizes the occurrence before it executes the next machine instruction and executes an I interrupt-handling microroutine or Procedure 602 instead of the next machine instruction. If the interrupt or 4 t fault cannot be handled by the Process 610 in which it occurs, the interrupt or fault results in a process swap. When the interrupt handling routine is finished, Process 610 which faulted or was interrupted can be returned to the CPU if it was removed and the next machine instruction executed.

S While the above method works well with faults, the fact that interrupts are asynchronous causes several problems: 114 Machine instructions cannot require an indefinite amount of time to execute, since W; -4L1~1- CI -512interrupts cannot be handled until the machine instruction during which they occur is finished.

It must be possible to remove a Process 610 from the CPU at any time, since the occurrence of an interrupt is not predictable. This requirement greatly increases the difficulty of process management.

The method used for interrupt and fault handling in a present embodiment of CS 10110 is described below.

2. Har dware-Tntetut-and-Paflt-f*adS I tO- S1 I 6 4' in-eS- 0110 In CS 10110, there are two levels of interrupts: those which may created and dealt with completely by software, and those which may created by hardware signals. The former class of interrupts is dealt with in the discussion of Processes 610; the latter, termed hardware interrupts, is discussed below.

In CS 10110, hardware interrupts and faults begin as invocations of microroutines in FU 10120. The invocations may be the result of Event signals or may be made by microprograms. For example, when IOS 10116 places data in MEM 10112 for JP 10114, an Inter-processor Message Event signal results, and the signal causes the invocation of Inter-processor Message Interrupt handler microcode. On the other hand, a Page Fault begins as an invocation of Page Fault microcode by LAT microcode. The actions taken by the microcode which begins handling the fault or interrupt depend on whether the fault or interrupt is handled by the Process 610 which was being -513executed when the fault or Event occurred or by a special KQS Process 610.

In the first case, the Event microcode may perform a Microcode-to-Software Call to a high-level language procedure which handles the Event. An example of an Event handled in this fashion is .a floating point overflow: when FU 10120 microcode determines that a floating point overflow has occurred, it invokes microcode which may invoke a floating point overflow procedure provided by the high-level language whose S- Language was being executed when the overflow occurred.

In alternate embodiments of CS 10110, the overflow procedure may also be in microcode.

In the second case, the microcode handling the fault or interrupt puts information in tables used by a KOS Process 610 which handles the fault or interrupt and then causes the KOS Process 610 to run at some later time by advancing an Event Counter awaited by the Process 610.

Event Counters and the operations on them are explained r 20 in detail in a following description of Processes 610.

Since the tables and Event Counters manipulated by 8 microcode are always present in MEM 10112, these operations do not cause page faults, and can be performed S 1, in monitor mode. For example, when IOS 10116 transmits an IPM Event signal to JP 10114 after IOS 10116 has loaded data into MEM 10112, the Event resulting from the Event signal invokes microcode which examines a queue containing messages from IOS 10116. The messages in the queue contain Event Counter locations, and the microcode ie L o r n d n, -514which examines the queue advances those Event counters, thereby causing Processes 610 which were waiting for the data returned by the I/O operation to recommence execution.

3. The--Montoar--Mdey----ifferentia- Mas-ing and- Hardware- ntrt. -Hfand -li FU 10120 micromachine's monitor mode and differential masking facilities allow a method of hardware interrupt handling which overcomes two problems associated with conventional hardware interrupt handling: an interrupt can be handled in a predictable amount of time regardless of the amount of time required to execute an SIN, and if the microcode which handles the interrupt executes in monitor mode, the interrupt may be handled at any time without unpredictable consequences. There are two sources of hardware interrupts in CS 10110: an Inter- Processor Message (IPM) and an Interval Timer 25410 Runout. An IPM occurs when IOS 10116 completes an I/O task for JP 10114 and signals completion of the task via TOJP Bus 10132. An Interval Timer Runout occurs when a preset time at which CS 10110 must take some action is reached. For: example, a given Process 610 may have a limit placed on the amount of time it may execute on JP 10114. As is explained in a following description of process synchronization, the virtual processor management system sets Interval Timer 25412 to run out when Process 610 has used all of the time available to it.

Both IPMs and Interval Timer Runouts begin as Event signals. The immediate effect of the Event signal is to set a bit in EP Field 27309 of MCW1. In principle, the Ii -515set bit can cause invocation of the event microcode for the Event on the next MO cycle in which the FU 10120 micromachine is in virtual mode. Since microroutines running in monitor mode are guaranteed to return the micromachine to virtual mode within a reasonable length of time, and the Event invocation will occur when this happens, the Event is guaranteed to be serviced in a reasonable period of time. The microroutines invoked by ti* the Events themselves execute in monitor mode, thereby 10 guaranteeing that no page faults will occur while they c are executing and that Process 610 which is executing on t L JP 10114 when the hardware interrupt occurs need not be removed from JP 10114.

While hardware interrupts are serviced in principle as described above, considerations of efficiency require that as many hardware interrupts as possible be serviced when the size of the FU micromachine's stack is at a minimum, at the beginning of an SIN's execution.

S This requirement is achieved by means of Egg TMR 25412 and ET Flag 27311 in MCW1 20290. As described above, when an IPM interrupt or an Interval Timer 25410 Runout interrupt occurs, Field 27317 or 27313 respectively is is S r set in MCW1 20290. At the same time, Egg TMR 25412 begins running. If the current SIN's execution ends before Egg TMR 25412 runs out, the set Field in MCW1 20290 causes the Interval Timer Runout or Inter-processor Message Event invocations to occur on the first microinstruction for the next SIN. If, on the other

I

-516hand, the current SIN's execution does not end before Egg TMR 25412 runs out, the Egg Timer Runout causes an Event signal. The immediate result of this signal is the setting of ST bit 27311 in MCW1 20290, and the setting of ET bit 27311 in turn causes the Interval Timer Runout Event invocation and/or IPM Event invocation to take place on the next MO cycle to occur while the micromachine is in virtual mode. The above mechanism thus guarantees that most hardware interrupts-will be 10 handled at the beginning of an SIN, but that hardware J interrupts will always be handled within a certain amount t l of time regardless of the length of time required to execute an SIN.

S g. PF Wcromachtine-and-Ce-ff--Stbsv sems The subsystems of CS 10110, such as the object subsystem, the process subsystem, the S-Interpreter subsystem, and the Name Interpreter subsystem, are implemented all or in part in the micromachine. The description of the micromachine therefore closes with an 20 overview of the relationship between these subsystems and the micromachine. Detailed descriptions of the operation of the subsystems have been presented previously.

The subsystems fall into three main groups: KOS subsystems, the 'ame Interpreter subsystem, and the S- Interpreter subsystem. The relationship between the three is to some extent hierarchical: the KOS subsystems provide the environment required by the Name Interpreter subsystem, and the Name Interpreter subsystem provides -517the environment required by the S-Interpreter subsystem.

For example, the S-Interpreter subsystem interprets SINs consisting of SOPs and Names; the Name Interpreter subsystem translates Names into logical descriptors, using values called ABPs to calculate the locations contained in the logical descriptors. The KOS subsystems calculate the values of the ABPs, translate Logical Descriptors 27116 into physical MEM 10112 addresses, and check whether a Process 610 has access to an object which it is referencing.

In a present embodiment of CS 10110, the Name Interpreter subsystem and the S-Interpreter subsystem are implemented completely in the micromachine; in other embodiments, they could be implemented in high-level languages or in hardware. The KOS subsystems are implemented in both the micromachine and in high-level language routines. In alternate embodiments of CS 10110, KOS subsystems may be embodied entirely in microcode, or in high-level language routines. Some high-level language routines may execute in any Process 610, while others are executed only by special KOS Processes 610.

i The KOS subsystems also differ from the others in the Smanner in which the user has access: with the S- Interpreter subsystem and the Name Interpreter subsystem, I the subsystems come into play only when SINs are executed; the subsystems are not directly visible to users of the system. Portions of the KOS subsystems, on the other hand, may be explicitly invoked in high-level S i -518language programs. For example, an invocation in a highlevel language program may cause KOS to bind a Process 610 to a Virtual Processor 612.

The following will first list the functions performed by the subsystems, and then relate the subsystems to the monitor and virtual micromachine modes and specific micromachine devices. KOS subsystems perform the following functions: Virtual memory management; Virtual processor management; Inter-processor communication; S' Access Control; C' U: Object management; and, ,I Process management.

The Name Interpreter performs the following functions: Fetching and parsing SOPs, and Interpreting Names.

S*f The S-Interpreter, finally, dispatches SOPs, i.e., locates the FU 10120 and EU 10122 microcode which executes the operation corresponding to a given SOP for a given S-Language.

Of these subsystems, the S-Interpreter, the Name Interpreter, and the microcode components of the KOS Pc process and object manager subsystems execute on the virtual micromachine; the microcode components of the remaining KOS subsystems execute on the monitor micromachine. As will be seen in the discussions of -519these subsystems, subsystems which execute on the virtual micromachine may cause Page Faults, and may therefore reference data located anywhere in memory; subsystems which execute on the monitor micromachine may not cause Page Faults, and the data bases which these subsystems manipulate must therefore always be present at known locations in MEM 10112.

The relationship between subsystems and FU 10120 micromachine devices is the following: Microcode for all S 10 subsystems uses DESP 20210, Microcode Addressing 27013, l and Register Addressing 27011, and may use EU Interface 27007. S-Interpreter microcode uses SOP Decoder 27003, I and Name Interpreter Microcode uses Instruction Stream Reader 27001, Parsing Unit 27005, and Name Translation Unit 27015. KOS virtual memory management microcode uses Memory Reference Unit 27017, and Protection Microcode uses Protection Unit 27019.

C* t Having described in detail the structure and operation of CS 10110's major subsystems, MEM 10112, FU S 20 10120, EU 10122, IOS 10116, and DP 10118, and the CS 10110 micromachine, CS 10110 operation will be described in further detail next below. First, operation of CS 10110's Namespace, S-Interpreter, and Pointer Systems will be described. Then, operation of CS 10110 will be described in further detail with respect to CS 10110's Kernel Operating System. It 4--,Ij~ f i"J -520- 3. Names pace, S-Inte-rr 307. 274) The preceding chapters have presented an overview of CS 10110, examined its hardware in detail, and explained how the FU 10120 hardware functions as a micromachine which controls the activities of other CS 10110 components. In the remaining portions of the specification, the means are presented by which certain key features of CS 10110 are implemented using the hardware, the micromachine, tables in memory, and highlevel language programs. The present chapter presents

C

s C three of these features: the Pointer Resolution System, INamespace, and the S-interpreters.

C The Pointer Resolution System translates pointers, data items which contain location information, into UID-offset addresses. Namespace has three main functions: Yt It locates SINs and fetches them from CS 10110's memory into FU 10120.

20 It parses SINs into SOPs and Names.

It translates Names into Logical Descriptors 1. 27116 or values.

The S-interpreters decode S-operations received from Snamespace into locations in microcode contained in FUSITT 11012 and EUSITT 20344 and then execute that microcode.

If the S-operations require operands, the S-interpreters use Namespace to translate the operands into Logical Descriptors 27116 or values as required by the operations.

-521- Since Namespace depends on the Pointer Resolution System and the S-interpreters depend on Namespace, the discussion of the systems begins with pointers and then deals with namespace and S-interpreters.

A. Pointers -and-Pointer-Reso-luton---Fis-.3---'30-2) A pointer is a data item which represents an address, in CS 10110, a UID-offset address. CS 10110 has I two large classes of pointers: resolved pointers and i c unresolved pointers. Resolved pointers are pointers 10 whose values may be immediately interpreted as UID-offset addresses; unresolved pointers are pointers whose values r must be interpreted by high level language routines or microcode routines to yield UID-offset addresses. The Sact of interpreting an unresolved pointer is called resolving it. Since the manner in which an unresolved pointer is resolved may be determined by a high-level language routine written by a system user, unresolved pointers provide a means by which users of the system may define their own pointer types.

Both resolved and unresolved pointers have subclasses. The subclasses of resolved pointers are UID pointers and object relative pointers. UID pointers "contain a UID and offset, and can thus represent any CS 10110 address; object-relative pointers contain only an offset; the address's UID is assumed to be the same as i that of the object containing the object-relative pointer. An object-relative pointer can therefore only represent addresses in the object which contains the pointer.

i i, L -522- The subclasses of unresolved pointers are ordinary unresolved pointers and associative pointers. The difference between the two kinds of unresolved pointers is the manner in which they are resolved. Ordinary unresolved pointers are always resolved by high-level language routines, while associative pointers are resolved the first time they are used in a Process 610 and a domain by high-level language routines, but are subsequently resolved by means of a table called the C 10 Associated Address Table (AAT). This table is accessible to microcode, and associative pointers may therefore be more quickly resolved than ordinary unresolved pointers.

The following discussion will first explain the formats used by all CS 10110 pointers, and will then explain how pointers are processed in FU 10120.

S*a. Pointer-F F rmats'Pq--'<-l-- Figure 301 represents a CS 10110 pointer. The figure has two parts: a representation of General Pointer Format 30101, which gives an overview of the fields which appear in all CS 10110 pointers, and a detailed presentation of .Flags and Format Field 30105, which contains the o information by which the kinds of CS 10110 pointers are distinguished.

Turning first to General Pointer Format 30101, all CS 10110 pointers contain 128 bits and are divided into three main fields: Offset Field 30103 contains the offset portion of a UID-offset address in resolved pointers and in associative pointers; in other unresolved f i.- -523pointers, it may contain an offset from some point in an object or other information as defined by the user.

Flags and Format Field 30105 contains flags and format codes which distinguish between kinds of pointers. These flags and format codes are explained in detail below.

S- UID field 30115 contains a UID in UID pointers t 4' and in some associative pointers; in object- 10 relative pointers, and other associative pointers, its meaning is undefined, and in ordinary unresolved pointers, it may contain information as defined by the user.

Flags and Format Field 30105 contains four subfields: Fields 30107 and 30111 are reserved and must be set to 0.

NR Field 30109 indicates whether a pointer is c resolved or unresolved. In resolved pointers, the field is set to 0, and in unresolved pointers, it is set to 1.

Format Code Field 30113 indicates the kind of resolved or unresolved pointers. Format codes for the present embodiment are explained below.

The values of Format Code Field 30113 may range from 0 to 31. If Format Code Field 30113 has the value 0, the pointer is a null pointer, a pointer which neither directly nor indirectly indicates an address. The meanings of the other format codes depend on the value of NR Field 30109: -524- NR Field Value Format Code Value Meaning 0 1 UID pointer 0 2 Object-relative pointer 0 all other codes Illegal 5 1 1 UID associative pointer 1 2 Object-relative associative pointer 1 all other codes Ordinary unresolved pointer As indicated by the above table, the present embodiment has two kinds of associative pointer, UID associative pointers and object-relative associative pointers. Like Sa UID pointer, a UID associative pointer contains a UID S, and an offset, and like an object-relative pointer, an object-relative associative pointer contains an offset and takes the value of the UID from the object to which it belongs. .However, as will be explained in detail m later, the UID and offset which the associative pointers contain or represent are not used as addresses. Instead, the UID and offset are used as tags to locate entries in the AAT, which associates an associative pointer with a resolved pointer.

b. Pointer3sinP.i lOf20-4P4g-:-TO2> When a pointer is used as an address in FU 10120, the address information in the pointer must be translated -525into a Logical Descriptor 27116 consisting of an AON, an offset, and a length field of 0; when a Logical Descriptor 27116 in FU 10120 is used to form a pointer value in memory, the AON must be converted back to a UID. The first conversion is termed pointer-todescriptor conversion, and the second descriptor-topointer conversion. Both conversions are accomplished by microcodes executing in FU 10120.

What is involved in the translation depends on the 10 kind of pointer: if the pointer is a UID pointer, the UID must be translated into an AON; if the pointer is an r object-relative pointer, the AON required to fetch the pointer is the pointer's AON, so no translation is necessary. If the pointer is an unresolved pointer, it must first be translated into a resolved pointer and then into a Logical Descriptor 27116. If the pointer is associative, the translation to a resolved pointer may be I performed by means of the ATT.

In the present embodiment, when other FU 10120 microcode calls pointer-to-descriptor microcode, the calling microcode passes Logical Descriptor 27116 for the «location of the pointer which is to be translated as an argument to the pointer-to-description translation microcode. The pointer-to-descriptor microcode returns a Logical Descriptor 27116 produced from the value of the pointer at the location specified by Logical Descriptor 27116 which the pointer-to-descriptor microcode received as an argument.

The pointer-to-descriptor microcode first uses Logical Descriptor 27116 given it as an argument to fetch -526the value of the pointer's Offset Field 30103 from memory. It then saves Logical Descriptor 27116's offset in the output register belonging to OFFALU 20242 and places the value of the pointer's Offset Field 30103 in the offset field of Logical Descriptor 27116 which it received as an argument. The pointer-to-descriptor microcode then saves Logical Descriptor 27116 indicating the pointer's location by storing Logical Descriptor c" 27116's AON and offset (obtained from OFFALU 20242) in a 1 0 register in the GRF 10354 frame being used by the invocation of the pointer-to-descriptor microcode. Next, the microcode adds 40 to the offset stored in OFFALU 20242, thereby obtaining the address of NR Field 30109, and uses the address to *fetch and read NR Field 30109 and Format Code Field 30113. The course of further processing is determined by the values of these fields.

If NR Field 30109 indicates a resolved pointer, there are I four cases, as determined by the value of Format Code Field 30113: Format code field 0: The pointer is a null pointer, S Format code field 1: The pointer is a UID pointer.

Format code field 2: The pointer is an intraobject pointer.

Any other value of the format code field: The pointer is invalid.

In the first case, the microcode sets all fields of the argument to Of in the second, it fetches the value of tID -527- Field 30115 from memory and invokes LAR microcode (explained in the discussion of objects), which translates the UID to the AON associated with it. The AON is then loaded into the argument's AON field. In the third case, the AON of Logical Descriptor 27116 for the pointer's location and the pointer's AON are the same, so the argument already contains the translated pointer. In the fourth case, the microcode performs a call to a S,0c pointer fault-handling Procedure 602 which handles t invalid pointer faults, passing saved Logical Descriptor S27116 for the pointer as an argument. Procedure 602 which handles the fault must return a resolved pointer to the microcode, which then converts it to a Logical Descriptor 27116 as described above.

c. Descriptot-to--Po+nter-Converson Descriptor to pointer conversion is the reverse of pointer to descriptor conversion with resolved pointers.

The operation must be performed whenever a resolved pointer is moved from an FU 10120 register into MEM 10112. The operation takes two arguments: a Logical S Descriptor 27116 which specifies the address to which the 1 pointer is to be written, and a Logical Descriptor 27116 whose AON and offset fields specify the location contained in the pointer. There are two cases: intraobject pointers and UID pointers. Both kinds of pointers have values in Offset Field 30103, so the descriptor-topointer microcode first writes the second argument's offset to location specified by the first argument's -528- Logical Descriptor 27116. The next step is to determine whether the pointer is an intra-object pointer or a UID pointer. To do so, the microcode compares the arguments' AONs. If they are the same, the pointer points to a location in the object which contains it, and is therefore an intra-object pointer. Since UID Field 30115 of an intra-object pointer is meaningless, the only step remaininc for intra-object pointers is to set Flags and Format Field 30105 to the binary representation of 2, which sets all bits but bit 46 to 0, and thereby identifies the pointer as a resolved intra-object pointer.

With UID pointers, the descriptor-to-pointer microcode sets Flags and Format Field 30105 to 1, thereby identifying the pointer as a resolved UID pointer, and calls a KOS LAR microroutine (explained in detail in the discussion of objects) which converts the first argument's AON to a UID and places the result UID in the current frame. When the KOS AON to UID conversion microroutine returns, the descriptor-to-pointer microcode Ewh. writes the UID to the converted pointer's UID Field 30115.

B. Name Dace and-theiS-ncecpaes theFg m0s-' Namespace and the S-interpreter both interpret information contained in Procedure Objects 608.

Consequently, the discussion of these components o CS 10110 begins with an overview of those parts of tcedure Object 606 relevant to Namespace and the S-interpreters, a. Procedure- Obiect-- 606--Overview--Fig-3- Figure 303 represents those portions of Procedure Object 608. Fig. 303 expands information contained in Fig. 103; Fields which appear in both Figures have the number of Fig. 103. Portions of Procedure Object 608 Q which are not discussed here are dealt with later in the discussion of Calls and Returns. The most important part 10 of a Procedure Object 608 for these systems is Procedure Environment Descriptor (PED) 30303. A Procedure 602's PED 30303 contains the information required by Namespace and the S-interpreter to locate and parse Procedure 602's code and interpret its Names. A number of Procedures 602 in a Procedure Object 608 may are a PED 30303. As will gbe seen in tl discussion of Calls the fact thatha Procedure 602 shares a PED 30303 with the Procedure 602 that invo kes it affects the manner in which the Call is executed.

The fields of FED 30303 which are important to the Spresent discussion are three fields in Header 30304: K

S

t Field 30305, LN Field 30307, and SIP Field 30309 2 and I P three of the remaining fields: NTP Field 30311 SDPP S tField 30313, and PBP Field 30315.

a K Field 30305 indicates whether the Names in the SINs of Procedures 602 which share FED 30303 have 8, 12, or 16 bits.

S LN Field 30307 contains the Name wh-c h has the largest isdex of any in Procedure 602's Name Table 10350.

Table 10350.

-530- SIP Field 30309 is a UID pointer to the object which contains the S-interpreter for Procedure 602's S-Language.

NTP Field 30311 is an object-relative pointer to the beginning of Procedure 602's Name Table 10350.

SDPP Field 30313 is a pointer which is resolved to the location of static data used by n Procedures 602 to which PED 30303 belongs when t V t one of Procedures 602 is invoked by a given Process 610. The resolved pointer corresponding to SDPP 30313 is the SDP ABP.

PBP Field 30315 contains the PBP ABP for invocations of Procedures 602 to which PED 30303 belongs. The PBP ABP is used to calculate locations inside Procedure Object 608.

Other areas of interest in Procedure Object 608 are Literals 30301 and Static Data Prototype (SDPR) 30317.

Literals 30301 contains literal values, values in Procedure 602 which are known at compile time and will not change during program execution. SDPR 30317 may P scontain' any of the following: pointers to external Sroutines and to static data contained in other objects, information required to create static data for a non-resolved.

In the present embodiment, Binder Area 30323 is also important. Binder Area 30323 contains information which

L

-r Il--nrrr-- -531ri I fi Ir t I I LI *1 I I I allows unresolved pointers contained in Procedure Object 608 to be resolved. Unresolved pointers other than SDPP 30313 in Procedure Object 608 all contain locations in Binder Area 30323, and the specified location contains the information required to resolve the pointer.

Fig. 303 contains arrows showing the locations in Procedure Object 608 pointed to by NTP Field 30311, SDPP Field 30313, and PBP Field 30315. NTP Field 30311 points to the beginning of Name Tables 10350, and thus a Name's 10 Name Table Entry can be located by adding the Name's value to NTP Field 30311. PBP Field 30315 points to the beginning of Literals 30301, and consequently, the locations of Literals and the locations of SINs may be expressed as offsets from the value of PBP Field 30315.

SDPP Field 30313 points to the beginning of SDPR 30317.

As will be explained in detail in the discussion of Calls, when a Procedure 602 has static data, the SDP ABP is derived from SDPP Field 30313.

b. Namespace The Namespace component of CS 10110 locates SINs belonging to a procedure.and fetches them from memory to FU 10120, parses SINs into SOPs and Names, and performs Resolve and Evaluation operations on Names. The Resolve operation translates a Name into a Logical Descriptor 27116 for the data represented by the Name, while the Evaluation operation obtains the data itself. The Evaluation operation does so by performing a Resolve operation and then using the resulting Logical Descriptor 'r j I -53?2- 27116 to fetch the data. Since the Evaluation and Resolve operations are the most complicated, the discussion begins with them.

1. Name Resoluton-ad--Evalu t-on Name Resolution and Evaluation translate Names into Logical Descriptors 27116 by means of information contained in the Names' NTEs, and the NTEs define locations in terms of Architectural Base Registers.

Consequently, the following discussion will first 10 describe Name Table Entries and Architectural Base Pointers and then the means by which Namespace translates the information contained in the Name Table Entries and Architectural Base Pointers into Logical Descriptors '27116.

2. The, Name- Tab-le---FI~-.--314-1 As previously mentioned, Name Tables 10350 are contained in Procedure Objects 608. Name Tables 10350 contain the information required to translate Names into t Logical Descriptors 27116 for the operands represented by 20 Lhe Names. Each Name has as its value the number of a Name Table Entry. A Name's Name Table Entry is located by multiplying the Name's value by the size of a short Name Table Entry and adding the product to the value in C ENTP Field 30311 of PED 30303 belonging to Procedure 602 which contains the SIN.

The Name Table Entry contains length and type information for the data item specified by the Name, and represents the data item's location as a displacement -533from a known location, termed the base. The base may be a location specified by an ABP, a location specified by another Name, or a location specified by a pointer. In the latter case, the pointer's location may be specified in terms of an ABP or as a Name.

Fig. 304 is a detailed representation of a Name Table Entry (NTE) 30401. There are two kinds of NTEs 30401: Short NTEs 30403 and Long NTEs 30405. Short NTEs 30403 contain 64 bits; Long NTEs 30405 contain 128 bits. Names 1 0 that represent scaler data items whose displacements may be expressed in 16 bits have Short NTEs 30403; Names that 1# represent scaler data items whose displacements require more than 16 bits and Names that represent array elements lave Long NTEs 30405.

A Short NTE 30403 has four main fields, each 16 bits in length: Flags and Format Field 30407 contains flags and format information which specify how Namespace is to interpret NTE 30401.

SV 20 Base Field 30425 indicates the base to which the displacement is to be added to obtain the location of the data represented by the Name.

Base Field 30425 may represent the location in i four ways: by means of an ABP, by means of a I Name, by means of a pointer located by means of r: I~an ABP, and by means of a pointer located by means of a Name.

i r -534- Length Field 30435 represents the length of the data. The length may be a literal value or a Name. If it is a Name, the Name resolves to a location which contains the data item's length.

Displacement Field 30437 contains the displacement of the beginning of the data from the base specified in Field 30425. The displacement is a signed integer value.

Long NTEs 30405 have four additional fields, each 16 bits 10 long? Two of the fields, Index Name Field 30441 and IES t_ Field 30445 are used only in NTEs 30401 for Names that represent arrays.

it t Displacement Extension Field 30439 is used in all Long NTEs 30405. If the displacement value in Field 30437 has less than 16 bits, Displacement Extension Field 30439 contains sign bits, the bits in the field are set to 0 when the displacement is positive and 1 when the t t displacement is negative. When the displacement S 20 value has more than 16 bits, Displacement S" Extension Fi'T3'b439 contains the most significant bits of the displacement value as well as sign bits.

Index Name Field 30441 contains a Name that represents a value used to index an element of an array.

Field 30443 is reserved.

IES Field 30445 contains a Name or Literal that specifies the size of an element in an array.

-535- The value represented by this field is used together with the value represented by Index Name Field 30441 to locate an element of an array.

As may be seen from the above, the following fields may contain names: Base Field 30425, Length Field 30435, Index Name Field 30441, and IES Field 30445.

Two fields in NTE 30401 require further consideration: Flags and Format Field 30407 and Base Field 30425. Flags and Format Field 30407 has three subfields: Flags Field 30408, FM Field 30421, and Type Field 30423. Turning first to Flag's Field 30408, the six flags in the field indicate how Namespace is to interpret NTE 30401. The flags have the following meanings when they are set: Long NTE Flag 30409: NTE 30401 is a Long NTE 30405.

Length is a Name Flag 30411: Length Field 30435 contains a Name.

Base is a Name Flag 30413: Base Field 30425 contains a Name instead of the number of an ABP.

Base Indirect Flag 30415: Base Field 30425 represents a pointer, and the location represented by NTE 30401 is to be calculated by obtaining the pointer's value and adding the value contained in Displacement Field 30437 and Displacement Extension Field 30439 to the pointer's offset.

Array Flag 30417: NTE 30401 represents an array.

-536- IES is a Name Flag 30419: IES Field 30445 contains a Name that represents the IES value.

Several of these flags may be set in a given NTE 30401.

For example, an entry for an array element that was referenced via a pointer to the array which in turn was represented by a Name, and whose IES value was represented by a Name, would have Flags 30409, 30413, 30415, 30417, and 30419 set.

FM Field 30421 indicates how the data represented by 10 the Name is to be formated when it is fetched from St Imemory. The value of FM Field 30421 is placed in FIU ,Field 27107 of Logical Descriptor 27116 produced from NTE 30401. The two bits allow for four possibilities: Setting Meaning S 15 00 right justify, zero fill 01 right justify, sign fill left justify, zero fill S. 11 left justify, ASCII space fill The four bits in Type Field 30423 are used by compilers S 20 for language-specific type information. The value of Type Field 30423 is placed in Type Field 27109 of Logical Descriptor 27116 produced from NTE 30401.

Base Field 30425 may have either Base is an ABP Format 30427 or Base is a Name Format 30432. The manner in which Base Field 30425 is interpreted depends on the setting of Base is a Name Flag 30413 and Base Indirect Flag 30415. There are four possibilities: I'.'jjl S'-537- -537- Field Settings Base is a Name Base Indirec 0 0 t Meaning ABP Format locates base directly.

ABP Format locates a pointer which is the base.

Base is Name Format locates base when Name is resolved.

Base is Name Format locates tttt i 1$ I I I i 6C i 61 I ti 10 a pointer when Name is resolve and the pointer is the base.

As indicated by the above table, Base Field 30425 is interpreted as having Base is ABP Format 30427 when Base is a Name Flag 30411 is not set. In Base is ABP Format 30427, Base Field 30425 has two subfields: ABP Field 30429 and Pointer Locator Field 30431. The latter field has meaning only when Base Indirect Flag 30415 is set.

ABP Field 30429 is a two-bit code which indicates the ABP. The settings and their meanings are the following: Setting APB 00 FP 01 Unused 10 SDP 11 PBP The ABPs are discussed below. When Base Indirect Flag 30415 is set to 1 and Base is a Name Flag 30413 is set to 0, the remaining 14 bits of the Base Field in ABP Format are interpreted as Pointer Locator Field 30413. When so 4 4.c i, 4 4 444144 I i -538interpreted, Pointer Locator Field 30413 contains a signed integer, which, when multiplied by 128, gives the displacement of a pointer from the ABP specified in ABP Field 30429. The value of this pointer is then the base to which the displacement is added.

Base Field 30425 is interpreted as having Base is a Name Format 30432 when Base is a Name Flag 30413 is set to 1. In Base is a Name Format 30432, Base Field 30425 contains a Name. If Base Indirect Flag 30415 is not set, 10 the Name is resolved to obtain the Base. If Base Indirect Flag 30415 is set, the name is evaluated to obtain a pointer value, and that pointer value is the Base.

3. Archt ectUral--Base- Po-nte eTS-f PiS2.

305- :306-) If Base is a Name Flag 30413 belonging to a NTE 30401 is not set, Base Field 30425 specifies one of the three ABPs in CS 10110: 1« PBP specifies a location in Procedure Object 608 to which displacements may be added to obtain S the locations of Literals and SINs.

SDP specifies a location in a Static Data Block for an invocation of a Procedure 602 to which 4 displacements may be added to obtain the locations of static data and linkage pointers to Procedures 602 contained in other Procedure Objects 608 and static data.

FP specifies a location in the MAS frame belonging to Procedure 602's current invocation pp- -539- 94 4 9 4 a 99 9 444 4 4 O *44499 9 9 94 99 49,.

to which displacements may be added to obtain the location of local data and linkage pointers to arguments.

Each time a Process 610 invokes a Procedure 602, Call microcode saves the current values of the ABPs on Secure Stack 10336, calculates the values of the ABPs for the new invocation, and places the resulting Logical Descriptors 27116 in FU 10120 registers, where they are accessible to Namespace microcode.

10 Call microcode calculates the ABPs as follows: PBP is obtained directly from PBP Field 30315 in PED 30303 belonging to the Procedure 602 being executed. All that is required to make it into a Logical Descriptor 27116 is the addition of the AON for Procedure Object 608's UID.

SDP is obtained by performing a pointer-to-descriptor translation on SDPP Field 30313. FP, finally, is provided by the portion of Call microcode which creates the new MAS 502 frame for the invocation. As is described in detail in the discussion of Call, the Call microcode copies linkage pointers to the invocation's actual arguments onto MAS 502, sets FP to point to the location following the last actual argument, and then allocates storage for the invocation's local data.

Positive displacements from FP thus specify locations in the local data, while negative offsets specify linkage pointers.

i r ni Ii i r i 9 99 99 t 9 9 t

-J

-540a.a. Reso-lving--and--Eva-htting--Names -(Fig0.- 3051 The primary operations performed by Namespace are resolving names and evaluating them. A Name has been resolved when Namespace has used the ABPs and information contained in the Name's NTE 30401 to produce a Logical Descriptor 27116 for the Name; a name has been evaluated when Namespace has resolved the Name, presented the resulting Logical Descriptor 27116 for the Name to 'tt* 10 memory, and obtained the value of the data represented by r t l the Name from memory.

S The resolve operation has three parts, which may be performed in any order: Obtaining the Base from Base Field 30425 of the Name's NTE 30401.

Obtaining the displacement.

j Obtaining the length from Length Field 30435.

Obtaining the length is the simplest of the operations: if Length in a Name Flag 30411 is set, the length is the v 20 value obtained by evaluating the Name contained.in Length Field 30435; otherwise, Length Field 30435 contains a literal value and the length is that literal's value.

O:tt There are four ways in which the Base may be I calculated. Which is used depends on the settings of Base is a Name Flag 30413 and Base Indirect Flag 30415: Both Flags 0: the ABP specified in ABP Field 30429 is the Base.

L

J

-541- Base is a Name Flag 30413 0 and Base Indirect Flag 30415 1: The Base is the location contained in the pointer specified by ABP Field 30429 and pointer Locator Field 30431.

Base is a Name Flag 30413 1 and Base Indirect Flag 30415 0: The Base is the location obtained by resolving the Name in Base Field 30425.

Both Flags 1: The Base is the location obtained t Sby evaluating the Name in Base Field 30425.

10 The manner in which Namespace calculates the Spdisplacement depends on whether NTE 30401 represents a Sscaler data item or an array data item. In the first case, Namespace adds the value contained in Displacement Field 30437 and Displacement Extension Field 30439 to the location obtained for the Base; in the second case, Namespace evaluates Index Name Field 30441 and IES Field 30445, multiplies the resulting values together, and adds Sthe product to the value in Displacement Field 30437 in order to obtain the displacement.

20 If any field of a NTE 30401 contains a Name, Namespace obtains the value or location represented by the Name by performing a Resolve or Evaluation operation on it as required. As mentioned in the discussion of NTEs 30401, flags in Flags Field 30408 indicate which fields of an NTE 30401 contain Names. Since the NTE 30401 for a Name used in another NTE 30401 may itself contain Names, Namespace performs the Resolve and Evaluation operations recursively.

r -542b.b. -Impementation--of--Name--gvaeluat-on and- Name--Reso-lve--in--eS---,O0 In the present embodiment, the Name Evaluation and Resolve operations are carried out by FU 10120 microcode Eval and Resolve commands. Both commands require two pieces of information: a register in the current frame of SR portion 10362 of GRF 10354 for receiving Logical Descriptor 27116 produced by the operation, and the source of the Name which is to be resolved or evaluated.

10 Both Resolve and Eval may choose between three sources: fI% St Parser 20264, Name Trap 20254, and the low-order 16 bits of the output register for OFFALU 20242. Resolve may ",4t specify current frame registers 0, 1, or 2 for Logical Descriptor 27116, and Eval may specify current frame registers 0 or 1. At the end of thd Resolve operation, Logical Descriptor 27116 for the data represented by the Name is in the specified SR 10362 register and at the end of the Evaluation operation, Logical Descriptor 27116 is in the specified SR 10362 register and the data's value 4 A 20 has been transferred via MOD Bus 10114 to EU 10122's OPB 20322.

The execution of both Resolve and Eval commands I always begin with the presentation of the Name to Name Cache 10226. The Name presented to Name Cache 10226 is latched into Name Trap 20254, where it is available for subsequent use by Name Resolve microcode.

If there is an entry for the Name in Name Cache 10226, a name cache hit occurs. For Names with NTEs 30401 fulfilling three conditions, the Name Cache 10226 entry for the Name is a Logical Descriptor 27116 for the p. -543data item represented by the Name. The conditions are the following: NTE 30401 contains no Names.

Length Field of NTE 30401 specifies a length of less than 256 bits.

If Base is Indirect Flag 30415 is set, Pointer Displacement Field 30431 must have a negative value, indicating that the base is a linkage pointer.

10 Logical Descriptor 27116 can be encached in this case r It \because neither the location nor the length of the data represented by the Name can change during the life of an r invocation of Procedure 602 to which the Name belongs.

If the Name Cache 10226 entry for the Name is a Logical Descriptor 27116, the hit causes Name Cache 10226 to place Logical Descriptor 27116 in the specified SR 10362 S, register. In all other cases, the Name Cache 10226 entry for the Name does not contain a Logical Descriptor 27116, and a hit causes Name Cache 10226 to emit a JAM signal.

S*C 20 The JAM signal invokes microcode which uses information stored in Name Cache 10226 to construct Logical Descriptor 27116 for the data item represented by the SName. JAMS are explained in detail below.

crt KIf there is no entry for the Name in Name Cache I 10226, a Name Cache Miss occurs, and Name Cache 10226 emits a cache miss JAM signal. The Name Resolve microroutine invoked by the cache miss JAM signal constructs an entry in Name Cache 10226 from the Name's

L.

-544- NTE 30401, using FU 10120's DESP 20210 to perform the necessary calculations. When it is finished, the cache miss microcode leaves a Logical Descriptor 27116 for the Name in the specified SR 10362 register and returns.

The Resolve operation is over when Logical Descriptor 27116 has been placed in the specified GRF 10354 register; the Evaluation operation continues by presenting Logical Descriptor 27116 to Memory Reference Unit 27017, which reads the data represented by Logical S' 10 Descriptor 27116 from memory and places it on OPB 20322.

The memory reference may result in Protection Cache 10234 S" misses and ATU 10228 misses, as well as protection faults and page faults, but these are handled by means of event signals and are therefore invisible to the Evaluation operation.

Name Cache 10226 produces 15 different JAM signals.

The signal produced by a JAM depends on the following: t whether the operation is a Resolve or an Eval, which register-Logical Descriptor 27116 is to be placed in, 20 whether a miss occurred, and in the case of a hit, which register in the Name Cache 10226 entry for the Name was loaded last. From the point of view of the behavior of the microcode invoked by the JAM, the last two factors 2 are the most important. Their relation to the microcode is explained in detail below.

In the present embodiment, all entries in Name Cache 10226 are invalidated when a Procedure 602 calls another Procedure 602. The invalidation is required because I

I

-545- Calls always change the value of FP and may also change the values of SDP and PBP, thereby changing the meaning of NTEs 30401 using displacements from ABPs. Entries for Names in invoked Procedure 602 are created and loaded into Name Cache 10226 when the Names are evaluated or resolved and a cache miss occurs.

The following discussion will first present Name Cache 10226 as it appears to the microprogrammer and then explain in detail how Name Cache 10226 is used to evaluate and resolve Names, how it is loaded, and how it is flushed.

c c.c. Name-- each---3-2-2--inties--tFig-.

-306- SThe structure and the physical behavior of Name Cache 10226 was presented in the discussion of FU 10120 hardware; here, the logical structure of Name Cache 10226 Sentries as they appear to the microprogrammer is i' presented. To the microprogrammer, Name Cache 10226 Sappears as a device which, when presented a Name on NAME S 20 Bus 20224, always provides the microprogrammer with a Name Cache 10226 entry for the Name consisting of four registers. The microprogrammer may read from or write to S any one of the four registers. When the microprogrammer t writes to the four registers, the action taken by Name Cache 10226 when a hit occurs on the Name associated with the four registers depends on which of the registers has most recently been loaded. The means by which Name Cache 10226 associates a Name with the four registers, and the

I

-546means by which Name Cache 10226 provides registers when it is full are invisible to the microprogrammer.

Fig. 306 illustrates Name Cache Entry 30601 for a Name. The four Registers 30602 in Name Cache Entry 30601 are numbered 0 through 3, and each Register 30602 has an AON, offset, and length field like those in GRF 10354 registers, except that some flag bits in GRF 10354 register AON fields are not included in Register 30602 te fields, and the length field in Register 30602 is 8 bits long. As is the case with GRF 10354 registers, the tr microprogrammer can read or write individual fields of t. Register 30602 or entire Register 30602. Name Cache Entry 30601 is connected via DB 27021 to DESP 20210, and consequently, the contents of a GRF 10354 register may be obtained from or transferred to a Register 30602 or viceversa. When the contents of a Register 30602 have been transfered to a GRF 10354 register, the contents may be processed using OFFALU 20242 and other arithmetic-logical devices in.DESP 20210.

20 d.d. Name Cache-10-2-26--Hffts When a Name is presented to Name Cache 10226 and Name Cache 10226 has a Name Cache Entry 30601 containing t information about the Name, a name cache hit occurs. On C a hit, Name Cache 10226 hardware always loads the contents of Register 30602 0 of the Name's Name Cache Entry 30601 into the GRF 10354 register specified in the Resolve or Eval microcommand. In addition, a hit may result in the invocation of microcode via a JAM: j r

A

I

-547f 1141 ft 4~ 4 4 5 5* S St S .1 1(1 1 Sf 55?

C

S N? S S E~ S S .5 S IS S 5551% S I The JAM may invoke special microcode for resolving Names of array elements whose NTEs 30401 allow certain hardware accelerations of indexc calculations.

The JAM may invoke general name resolution microcode which produces a Logical Descriptor 27116 from the contents of Name Cache Entry 30601.

Whether the hit produces a JAM, and the kind of JAM it t0 produces, are determined by the last Register 30602 to be loaded when Name Cache Entry 30601 was created by Name Cache Miss microcode. If Register 30602 0 was the last to be loaded, no JAM occurs; if Register 30602 1 was loaded last,' the JAM for special array Name resolution occurs; if Register 30602 2 or 3 was loaded last, the JAM for general Name resolution occurs.

As may be inferred from the above, Name Cache 10226 hardware defines the manner in which Name Cache Entries 30601 are loaded for the first two cases. In the first 20 case, Name Cache Register 30602 0 must contain Logical Descriptor 27116 for the Name's data. As already mentioned, the Name's NTE 30401 must therefore describe data whose location and length does not change during an invocation and whose length is less than 256 bits. Name Cache 10226 hardware also determines the form of Name Cache Entries 30601 for encachable arrays. An encachable array NTE 30401 is an array NTE 30401 which fills the following conditions: -548- The only Name contained in array NTE 30401 is in Index Name Field 30441.

S- NTE 30401 for the index Name fills the conditions for scaler NTEs 30401 for which Logical Descriptors 27116 may be encached.

The value in IES Field 30445 is no greater than 128 and a power of 2.

Array NTE 30401 otherwise fills the conditions for scaler NTEs 30401 for which Logical S 10 Descriptors 27116 may be encached.

S"In the present embodiment, the encachable array entry uses registers 0, 1, and 2 of Name Cache Entry 30601 for the name: Register Contents AON OFFSET LENGTH 0 Logical Descriptor 27116 for the index Name.

1 0 IES power of 2 unused 2 Logical Descriptor 27116 for the array 20 When a hit for this type of entry occurs, the resulting JAM signal does two things: it invokes encachable array resolve microcode and it causes the index Name's Logical Descriptor 27116 to be presented to Memory Reference Unit 27017 for a read operation which returns the value of the data represented by the index Name to an accumulator in OFFALU 20242. The encachable array resolve microroutine then uses the Name that caused the JAM, latched into Name Trap 20254, to locate Register 30602 2 of Name Cache -549- Entry 30601 for the Name, writes the contents of Register S30602 2 into the GRF register specified by the Resolve or Eval microcommand, obtains the product of the IES value and the index value by shifting the index value left the number of times specified by the IES exponent in Register 30602 1, adds the result to the offset field of the GRF 10354 register containing the array's Logical Descriptor 27116, thus obtaining Logical Descriptor 27116 for the desired array element, and returns.

S 10 For the other cases, the manner in which Name Cache Entries 30601 are loaded and processed to obtain Logical I t Descriptors 27116 is determined by the microprogrammer.

The JAM signal which results if a Name Cache Entry 30601 i is neither a Logical Descriptor 27116 nor an encachable array entry merely invokes a microroutine. The microroutine uses the Name latched into Name Trap 20254 S to locate the Name's Name Cache Entry 30601 and then reads tag values in Name Cache Entry 30601 to determine how the information in Name Cache Entry 30601 is to be S 20 translated into a Logical Descriptor 27116. The contents of Name Cache Entries 30601 for the other cases have two general forms: one for NTEs 30401 with Base is Indirect Flag 30415 set, and one for NTEs in which it is not set.

tThe first general form looks like this: Register Contents SAON OFFSET LENGTH 0 ABP AON tag length unused 1 0 index name/ IES unused -550- 2 0 unused unused 3 0 data displacement unused from loc. specified by pointer Register 30602 0 contains the AON of the ABP. Register 30602 O's offset field contains two items: the tag, which contains Flags Field 30408 of NTE 30401 along with other information, and which determines how Name Resolve #of microcode interprets the contents of Name Cache Entry 30601, and a value or Name for the length of the data item. Register 30602 1 is used only if the Name represents a data item in an array. It then contains the r Name from Index Field 30441 and the Name or value from IES Field 30445. The offset field of Register 30602 3 contains the sum of the offset indicated by NTE 30401's ABP and of the displacement indicated by NTE 30401.

The second format, used for NTEs 30401 whose bases are obtained from pointers or by resolving a Name, looks like this: Registers Contents, AON OFFSET

LENGTH

0 0 tag length unused 1 0 index name/ IES unused 2 0 FM and type bits/ otnused base field 3 0 data displacement unused from loc. specified by pointer or name -551- In this form, the location of the Base must be obtained either by evaluating a pointer or resolving a Name.

Hence, there is no field specifying the Base's AON.

Otherwise, Registers 30602 0 and 1 have the same contents as in the previous format. In Register 30602 2, the offset field contains Name Table Entry 30401's FM Field 30421 and Type Field 30423 and Base Field 30425. The .Offset Field of Register 30602 2 contains the value of Name Table Entry 30401 Displacement Fields 30437 and S, 10 30439.

I As in Name Table Entries 30401, the index must be represented by a Name, and length, IES, and Base may be .I represented by Names. If a field of Name Cache Entry S, 30601 contains a Name, a flag in the tag indicates that fact, and Name Resolve microcode performs an Eval or Resolve operation on it as required to obtain the value or location represented by the name.

SL. The microcode which resolves Name Cache Entries 30601 of the types just described uses the general algorithms 20 described in the discussion of Name Table Entries 30401, and is therefore not discussed further here.

e.e. Name--Cache--O0-26--M-sisses .When a Name is presented to Name Cache 10226 and i *.so there is no Name Cache Entry 30601 for the Nar;e, a name cache miss occurs. On a miss Name Cache 10226 hardware emits a JAM signal which invokes name cache miss microcode. The microcode obtains the Name which caused the miss from Name Trap 20254 and locates the Name's NTE ~c -552- 30401 by adding the Name to the value of NTP 30311 from PED 30303 for Procedure 602 being executed. As will be explained in detail later, when a Procedure 602 is called, the Call microcode places the AON and offset specifying the NTP's location in a register in GR's 10360. Using the information contained in the Name's NTE 30401, the Cache Miss microcode resolves the Name and constructs a Name Cache Entry 30601 for it. As described above, the microcode determines the method by which it 10 resolves the Name and the form of the Name's Name Cache (Entry 30601 by reading Flags Field 30408 in the Name's NTE 30401. Since the descriptions of the Resolve operation, the micromachine, Name Cache 10226, and the t formats of Name Cache Entries 30601 are sufficient to allow those skilled in the art to understand the operations performed by Cache Miss microcode, no further description of the microcode is provided.

F.lushing--Name-Cache---2-26 As described in the discussion of Name Cache 10226 S 20 hardware, hardware means, namely VALS 24068, exist which allow Name Cache Entries 30601 to be invalidated. Name Cache Entries 30601 may be invalidated singly, or all entries in Name Cache 10226 may be invalidated by means of a single microcommand. The latter operation is termed name cache flushing. In the present embodiment, Name Cache 10226 must be flushed when Process 610 whose Virtual Processor 612 is bound to JP 10114 executes a Call or a Return and whenever Virtual Processor 612NO is

.I

t .i i -553unbound from JP 10114. Flushing is required on Call and Return because Calls and Returns change the values of the ABPs and other pointers needed to resolve Names. At a minimum, a Call produces a new MAS Frame 10412, and a Return returns to a previous Frame 10412, thereby changing the value of FP. If the called Procedure 602 has a different PED 30303 from that of the calling Procedure 602, the Call or Return may also change PBP, SDP, and'NTP. Flushing is required when a Virtual 10 Processor 612 is unbound from JP 10114 because Virtual Processor 612 which is next bound to JP 10114 is bound to a different Process 610, and therefore cannot use any information belonging to Process 610 bound to the previous Virtual Processor 612.

g.g. Petch-ing-the--Stream As explained in the discussion of FU 10120 hardware, S* SINs are fetched from memory by Prefetcher 20260. PREF 20260 contains a Logical Descriptor 27116 for a location in Code 10344 belonging to Procedure 602 which is S 20 currently being executed. On any MO cycle, PREF 20260 can place Logical Descriptor 27116 on DB 27021, cause Memory Reference Unit 27017 to fetch 32 bits at the location specified by Logical Descriptor 27116, and write S them into INSTB 20262. When INSTB 20262 is full, PREF A 25 20260 stops fetching SINs until Namespace parsing operations, described below, have processed part of tho contents of INSTB 20262, thereby creating space for more SINs.

-554- The fetching operation is automatic, and requires intervention from Namespace only when a SIN causes a branch, causes the next SIN to be executed to be some other SIN than the one immediately following the current SIN. On a branch, Namespace must load PREF 20260 with the location of the next SIN to be executed and cause PREF 20260 to begin fetching SINs at that locatiO6. The operation which does this is specified by the load.prefetchfor_branch microcommand. The microcommand specifies a source for a Logical Descriptor 27116 and transfers that Logical Descriptor 27116 via DB 27021'to PREF 20260. After PREF 20260 has thus been loaded, it begins fetching SINs at the specified location. Since any SINs still in INSTB 20262 have been rendered meaningless by the branch operation, the first SINs loaded into INSTB 20262 are simply written over INSTB 20.62's prior contents. Fig. 274 contains an example of the use of the load.prefetchforbranch microcommand.

S 20 h.h. Parsing-the-- -Stream The I-stream as fetched from MEM 10112 and stored in INSTB 20262 is a sequence of SOPs and Names. As already :mentioned, the I-stream has a fixed format: in the r a *present embodiment, SOPs are always 8 bits long, and 1 25 Names may be 8, 12, or 16 bits long. The length of Names used in a given procedure is fixed, and is indicated by the value in K Field 30305 in the Procedure 602's PED 30303. The Namespace parsing operations obtain the SOPs .f -555-

I,

i f I Sr I I I and Names from the I-stream and place them on NAME Bus 20224. The SOPs are transferred via this bus to the devices in SOP Decoder 27003, while the Names are transferred to Name Trap 20254 and Name Cache 10226 for S Resolve and Evaluation operations as described above. As the parsing operations obtain SOPs and Names, they also update the three program counters CPC 20270, EPC 20274, and IPC 20272. The values in these three counters are offsets from PBP which point to locations in Code 10344 10 belonging to Procedure 602 being executed. CPC 20270 points to the I-stream syllable currently being parsed, so it is updated on every parsing operation. EPC 20274 points to the beginning of the last SIN executed by JP 10114, and IPC 20272 points to the beginning of the current SIN, so these program counters are changed only at the beginning of the execution of an SIN, when a SOP is parsed.

As described in the discussion of FU 10120 hardware, in the current implementation, parsing consists 20 physically of reading 8 or 16 bits of data from a location in INSTB 20262 identified by a pointer for INSTB 20262 which is accessible only to the hardware. As data is read, the hardware increments the pointer by the number of bits read, wrapping around and returning to the beginning of INSTB 20262 if it reaches the end. At the same time that the hardware increments the pointer, it increments CPC 20270 by the same number of bits. As previously mentioned, CPC 20270 contains the offset from I Il S $5 I C0 -556- PBP of the SOP or Name being currently parsed, thus coordinating the reading of INSTB 20262 with the reading of Procedure 602's Code 10344.

The number of bits read depends on whether Parser 20264 is reading an SOP or a Name, and in the latter case, by the syllable size specified for the Name. The syllable size is contained in CSSR 24112. On a Call to a Procedure 602 which has a,different PED 30303 from that of the calling procedure, the Call microcode loads the 10 value contained in K Field 30305 into CSSR 24112.

Namespace's parsing operations are performed by separate microcommands for parsing SOPs and Names. There is a single microcommand for parsing S-operations: parseopstage. The microcommand obtains the next eight 1 5 bits from INSTB 20262, places the bits onto NAME Bus 20224, and latches them into LOPCODE Register 24212. It St also updates EPC 20274 and IPC 20272 as required at the .beginning of an SIN: EPC 20274 is set to IPC 20272's former value, and IPC 20272 is set to CPC 20270's value.

S 20 At the end of the operation, CPC 20270 is incremented by 8. Since the parsing of an SOP always occurs as the first operation in the interpretation of an SIN, the parse-oplstage command is generally combined with a c dispatch fetch command. As will be explained below, the latter command interprets the S-operation as an address in FDISP 24218, and FDISP 24218 in turn produces an address in FUSITT 11012. The latter address is the location of the beginning of the SIN microcode for the i

SIN.

4 -557- 4t 4 4,i 4 4 4 t4 4 There are two microcommands for parsing Names: parsekloaduepc and parsekdispatchebox. Both commands obtain a number of bits from INSTB 20262 and place them on NAME Bus 20214. With both microcommands, the syllable size, stored in CSSR 24112, determines the number of bits obtained from INSTB 20262. Both commands also increment CPC by the value stored in CSSR 24112. In addition, parse_ k loadepc sets EPC to IPC's value, while parsekdispatch-ebox also dispatches EU 10122, interprets the SOP saved in LOPCODE 24210 as an address in EDISP 24222, which in turn contains an address in EU EUSITT 20344. The EU EUSITT 20344 address is passed via EUDIS Bus 20206 to COMQ 20342 in EU 10122.

c. The-S-nrtepreters-Pigr-3 CS 10110 does not assign fixed meanings to SOPs.

While all SOPs are 8 bits long, a given 8 bit SOP may have one meaning in one S-Language and a completely different meaning in another S-Language. The semantics of an S-Language's S-operations are determined completely 20 by the S-interpreter for the S-Language. Thus, in order to correctly interpret an S-operation, CS 10110 must know what S-interpreter it is to use. The S-interpreter is identified by a UID pointer with offset 0 in SIP Field 30309 of PED 30303 for Procedure 602 that CS 10110 is 25 currently executing. In the present embodiment, the UID is the UID of a microcode object which contains FU 10120 microcode. When loaded into FUSITT 11012, the microcode interprets SOPs as defined by the S-Language to which the 4c t 4 4*.

*r 4

'II

t i:l -558- SOP belongs. In other embodiments, the UID may be the UID of a Procedure Object 608 containing Procedures 602 which interpret the S-Language's SOPs, and in still others, the S-interpreter may be contained in a PROM and the S-interpreter UID may not specify an object, but may serve solely to identify the S-interpreter.

When a Procedure 602 executes an SIN on JP 10114, CS 10110 must translate the value of SIP Pointer 30309 for Procedure 602 and the S-instruction's SOP into a location *te 10 in the microcode or high-level language code which makes I up the S-interpreter. The location obtained by the translation is the beginning of the microcode or high- S, level language code which implements the SIN. The translation of an SOP together with SIP Pointer 30309 into a location in the S-interpreter is termed dispatching. Dispatching in the present embodiment involves two primary components: a table in memory which translates the value of SIP Pointer 30309 into a small integer called the Dialect Number, and S-operation Decoder Portion 27003 of the FU 10120 micromachine. The following discussion will first present the table and explain how an SIP Pointer 30309 is translated into a too Dialect Number, and then explain how the Dialect Number t 1 6 0 and the SOP together are translated into locations in FUSITT 11012 and EUSITT 20344.

I. TIansglatirne-S--intQ-a-p-Dazeg-Nmbe In the present embodiment, all S-interpreters in CS 10110 are loaded into FUSITT 11012 when CS 10110 begins

L..

-559operation and each S-interpreter is always placed in the same location. Which S-interpreter is used to interpret an S-Language is determined by a value stored in dialect register RDIAL 24212. Consequently, in the present embodiment, a Call to a Procedure 602 whose S-interpreter differs from that of the calling Procedure 602 must translate the UID pointer contained in SIP Field 30309 into a Dialect Number.

Fig. 307 represents the table and microcode which 10 performs this translation in the present embodiment. Sinterpreter Translation Table (STT) 30701 is a table St*, which is indexed by small AONs; Each STT Entry (STTE) 30703 has two fields: an AON Field 30705 and a Dialect Number Field 30709. Dialect Number Field 30709 contains the Dialect Number for the S-interpreter object whose AON is in AON Field 30705.

When CS 10110 begins operation, each S-interpreter object is wired active and assigned an AON small enough to serve as an index in STT 30701. By convention, a 20 given S-interpreter object is always assigned the same AON and the same Dialect Number. The AON is placed in AON Field 30705 of STTE 30703 indexed by the AON, and the Dialect Number is placed in Dialect Number Field 30709.

Since the S-interpreter objects are wired active, these AONs will never be reassigned to other objects.

On a Call which requires a new S-interpreter, Call microcode obtains the new SIP from SIP Field 30309, calls KOS LAR microcode to translate its UID to its AON, uses the AON to locate the S-interpreter's STTE 30703, and -560places the value of Dialect Number Field 30709 into RDIAL 21242, Other embodiments may allow S-interpreters to be loaded into FUSITT 11012 at times other than system initialization, and allow S-interpreters to occupy different locations in FUSITT 11012 at different times.

In these embodiments, STT 30701 may be implemented in a manner similar to the implementations of AST 10914 or MHT 10716 in the present embodiment.

It t 10 2. Dispatchina Dispatching is accomplished by Dispatch Files 27004.

These files translate the values provided by RDIAL 24212 and the SOP of the S-instruction being executed into the location of microcode for the SIN specified by the Soperation in the S-interpreter specified by the value of RDIAL 24212. The present embodiment has three dispatch S files: FDISP 24218, FALG 24220, and EDISP 24222. FDISP S24218 and FALG 24220 translate S-operations into locations of microcode which executes on FU 10120; EDISP S 20 24222 translates S-operations into locations of microcode which executes on EU 10122. The difference between FDISP 24218 and FALG 24220 is one of speed: FDISP 24218 can S* translate an SOP in the same microinstruction which performs a parse.op:stage command to load the SOP into LOPCODE 24210. FALG 24220 must perform the translation on a cycle following the one in which the SOP is loaded into LOPCODE 24210. Typically, the location of the first portion of the microcode to execute an S-operation is L 7 i- S-561cctained in an FDISP 24218 register, the location of portions executed later is contained in an FALG 24220 register, and the location of microcode for the Soperation which executes on EU 10122 is contained in EDISP 24222.

In the present embodiment, the registers accomplish the translation from S-operation to microcode location as follows: As mentioned in the discussion of FU 10120.

hardware, each Dispatch File contains 1024 registers.

S,,o 10 Each register may contain an address in an Sinterpreter. As will be seen in detail later, the address may be an address in an S-interpreter's object, or it may be the address in FUSITT 11012 or EUSITT 20344 of a copy of microcode stored at an S-interpreter address. The registers in the Dispatch Files may be divided into sets of 128 or 256 registers. Each set of S' registers translates the SOPs for a single S-Language into locations in microcode. Which set of registers is used to interpret a given S-operation is decided by the value of RDIAL 24212; which register in the set is used is determined by the value of the S-operation. The value contained in the specified register is then the location of microcode which executes the S-instruction specified ue t iby the S-operation in the S-Language specified by RDIAL 24212.

Logically, the register addressed by the concatenated value in turn contains a 15 bit address which is the location in the S-interpreter of the first -562microinstruction of microcode used to execute the Sinstruction specified by the S-operation in the S- Language specified by the contents of RDIAL 24212. In the present embodiment, the microcode referred to by the address may have been loaded into FUSITT 11012 and EUSITT 20344 or it may be available only in memory. Addresses of microcode located in FUSITT 11012 and EUSITT 20344 are only eight bits long. Consequently, if a Dispatch File 27004 contains an address which requires more bits than 10 that, the microcode specified by the address is in memory. As described in the discussion of FU 10112 hardware, addresses larger than 8 bits produce an Event Signal, and microcode invoked by the event signal fetches the microinstruction at the specified address in the Sinterpreter from memory and loads it into location 0 of FUSITT 11012. The event microcode then returns, and the microinstruction at location 0 is executed. If the next microinstruction also has an address larger than 8 bits, the event signal occurs again and the procesSs described o 20 above is repeated.

As previously mentioned, FDISP 24218 is faster than FALG 24220. The reason for the difference in spesd is that FDISP registers contain only 6 bits for addressing .the S-interpreter. The present embodiment assumes that all microcode addressed via FDISP 24218 is contained in SFUSITT 11012. It concatenates 2 zero bits with the six bits in the FDISP 24218 register to produce an 8 bit address for FUSITT 11012. FDISP 24218 registers can thus 7 -563contain the location of every fourth FUSITT 11012 register between FUSITT register 256 and FUSITT register 448. The microcode loaded into these locations in FUSITT 11012 is microcode for operations which are performed at the start of the SIN by many different SINs. For example, all SINs which perform operations on 2 operands and assign the result to a location specified by a third operand must parse and evaluate the first two operands and parse and resolve the third operand. Only after l 10 these operations are done are SINs-specific operations f 4 performed. In the present embodiment, the microcode t which parses, resolves, and evaluates the operands is contained in a part of FUSITT 11012 which is addressable by FDISP 24218.

As previously mentioned, in the present embodiment, FUSITT 11012 and EUSITT 20344 may be loaded only when CS 10110 is initialized. The microcode loaded into FUSITT 11012 and EUSITT 20344 is produced by the microbinder from the microcode for the various SINs. To achieve 20 efficient use of FUSITT 11012 and EUSITT 20344, microcode for operations shared by various S-interpreters appears only once in FUSITT 11012 and EUSITT 20344. While the t SINs in different S-Languages which share the microcode S. have different registers in FDISP 24218, FALG 24220, or EDISP 24222 as the case may be, the registers for each of the S-instructions contain the same location in FUSITT 11012 or EUSITT 20344.

-564- 4. The Kernel Operating System A. Introduction Many of the unique properties of CS 10110 are produced by the manipulation of tables in MEM 10112 and Secondary Storage 10124 by programs executing on JP 10114. These programs and tables together make up the Kernel Operating System (KOS). Having described CS 10110's components and the means by which they cooperate to execute computer programs, this specification now 10 presents a detailed account of KOS and of the prope-ties 0* of CS 10110 which it produces. The discussion begins *with a general introduction to operating systems, then presents an overview of CS 10110's operating systems, an overview of the KOS, and detailed discussions of the 15 implementation of objects, access control, and Processes 610.

C a. Oeratn.stems °o In CS 10110, as in other computer systems, the operating system has two functions: 20 It controls the use of CS 10110 resources such as JP 10114, MEM 10112, and devices in IOS 10116 by programs being executed on CS 10110.

It defines how CS 10110 resources appear to users of CS 10110.

The second function is a consequence of the first: By controlling the manner in which executing programs use system resources, the operating system in fact determines how the system appears to its users. Figure 401 is a -565schematic representation of the relationship between User 40101, Operating System 40102, and System Resources 40103. When User 40101 wishes to use a System Resource 40103, User 40101 requests the use of System Resource 40103 from Operating System 40102, and Operating System 40102 in turn commands CS 10110 to provide the requested Resources 40103. For example, when a user program wishes to use a peripheral device, it does not deal with the device directly, but instead calls the Operating System 10 40102 Procedure 602 that controls the device. While ft** Operating System 40102 must take into account the V device's complicated physical properties, the user S. program that requested the device need know nothing about the physical properties, but must only know what information the Operating System 40102 Procedure 602 requires to perform the operation requested by the user E't program. For example, while the peripheral device may require that a precise pattern of data be presented to it, the Operating System 40102 Procedure 602 may only require the data itself from the user program, and may format the data as required by the peripheral device.

The Operating System 40102 Procedure 602 that controls the peripheral device thus transforms a complicated 1"84 .physical interface to the device into a much simpler logical interface.

1. Reser ces--Con4-r^ed--bv-epe0a^:ng Syntems O2" t Operating Systems 40102 control two kinds of u resources: physical resources and virtual resources. The -566physical resources in the present embodiment of CS 10110 are JP 10114, IOS 10116 and the peripheral devices associated with IOS 10116, MEM 10112, and Secondary Storage 10124. Virtual resources are resources that the operating system itself defines for users of CS 10110.

As was explained above, in controlling how CS 10110's resources are used, Operating System 40102 defines how CS 10110 appears to the users. Instead of the physical resources controlled by Operating System 40102, the user St. 10 sees a far simpler set of virtual resources. The logical I/O device interface that Operating System 40102 gives r: the user of a physical I/O device is such a virtual Scresource. Often, an Operating System 40102 will define sets of virtual resources and multiplex the physical resources among these virtual resources. For instance, Operating System 40102 may define a set of Virtual Processors 612 that correspond to a smaller group of physical processors, and a set of virtual memories that ,t correspond to a smaller group of physical resources.

When a user executes a program, it runs on a Virtual Processor 612 and uses virtual memory. It seems to the user of the virtual processor and the virtual memory that he has sole access to a physical processor and physical s memory, but in fact, Operating System 40102 is multiplexing the physical processors and memories among the Virtual Processors 612 and virtual memories.

Operating System 40102, too, uses virtual resources.

For instance, the memory management portion of an -567- Operating System 40102 may use I/O devices; when it does so, it uses the virtual I/O devices defined by the portion of the Operating System 40102 that manages the I/O devices. One part of Operating System 40102 may also redefine virtual resources defined by other parts of Operating System 40102. For instance, one part of Operating System 40102 may define a set of primitive virtual I/O devices and another part may use these primitive virtual I/O devices to define a set of high- 10 level user-oriented I/O devices. Operating System 40102 .thus turns the physical CS 10110 into a hierarchy of virtual resources. How a user of CS 10110 perceives CS 10110 depends entirely on the level at which he is dealing with the virtual resources.

The entity that uses the resources defined by Operating System 40102 is the Process. A Process 610 may be defined as the activity resulting from the execution 5 of a program with its data by a sequential processor.

Whenever a user requests the execution of a program on CS 10110, Operating System 40102 creates a Process 610 which then executes the Procedures 602 making up the user's o program. In physical terms, a Process 610 is a set of data bases in memory that contain the current state of the program execution that the process represents.

Operating System 40102 causes Process 610 to execute the program by giving Process 610 access to the virtual resources which it requires to execute the program, by giving the virtual resources access to those parts of -568- Process 610's state which they require to perform their operations, and by giving these virtual resources access to the physical resources. The temporary relationship of one resource to another or of a Process 610 to a resource is called a binding. When a Process 610 has access to a given Virtual Processor 612 and Virtual Processor 612 has access to Process 610's state, Process 610 is bound to Virtual Processor 612, and when Virtual Processor 512 has access to JP 10114 and Virtual Processor 612's state is 22 *10 loaded into JP 10114 registers, Virtual Processor 612 is bound to JP 10114, and JP 10114 can execute SINs contained in Procedures 602 in the program being executed by Process 610 bound to Virtual Processor 612. Binding and unbinding may occur many times in the course of the execution of a program by a Process 610. For instance, if a Process 610 executes a reference to data and the data is not present in MEM 10112, then Operating System 40102 unbinds Process 610's Virtual Processor 512 from JP 10114 until the data is available in MEM 10112. If the data is not available for an extended period of time, or if the user for whom Process 610 is executing the program wishes to stop the execution of the program for a while, Operating System 40102 may unbind Process 610 from its Virtual Processor 612. Virtual Processor 612 is then available for use by ether Processes 610.

As mentioned above, the binding process involves giving a first resolro- access to a second resource, and using the first resource's state in the second resource.

To permit binding and unbinding, Operating System 40102 -569maintains data bases that contain the current state of each resource and each Process 610. State may be defined as the information that the operating system must have to use the resource or execute the Process 610. The state of a line printer, for instance, may be variables ),hat indicate whether the line printer is busy, free, off line, or out of order. A Process 610's state is more involved, since it must contain enough information to #to allow Operating System 40102 to bind Process 610 to a 1 ]0 Virtual Processor 612, execute Process 610 for a while, unbind Process 610, and then rebind it and continue 'r execution where it was halted. A Process 610's state thus includes all of the data used by Process 610 up to the time that it was unbound from a Virtual Processor 612, along with information indicating whether Process 610 is ready to begin executing again.

SFigure 402 shows the relationship between Processes 610, virtual, and physical resources in an operating system. The figure shows a multi-process Operating S" 20 System 40102, that is, one that can multiplex CS 10110 resources among several Processes 610. The Processes 610 thus appear to be executing concurrently. The solid arrows in Figure 402 indicate bindings between virtual *t 1 *resources or between virtual and physical resources.

Each Process 610 is created by Operating System 40102 to execute a user program. The program consists of Procedures 602, and Process 610 executes Procedures 602 in the order prescribed by the program. Processes 610 are created and managed by a component of Operating System 40102 called the Process Manager. Process Manager 40203 -570executes a Process 610 by binding it to a Virtual Processor 612. There may be more Processes 610 than there are Virtual Processors 612. In this case, Operating System 40102 multiplexes Virtual Processors 612 among Processes 610.

Virtual Processors 612 are created and made available by another component of Operating System 40102, Virtual Processor Manager 40205. Virtual Processor Manager 40205 also multiplexes JP 10114 among Virtual Processors 612.

If a Virtual Processor 612 is ready to run, Virtual Processor Manager 40205 binds it to JP 10114. When "Virtual Processor 612 can run no longer, or when another a Virtual Processor 612 requires JP 10114, Virtual Processor Manager 40205 unbinds running Virtual Processor 612 from JP 10114 and binds another Virtual Processor 612 to it.

Virtual Processors 612 use virtual memory and I/O resources to perform memory access and input-output.

Virtual Memory 40206 is created and managed by Virtual Memory Manager 40207, and Virtual I/O Devices 40208 are created and managed by Virtual I/O Manager 40209. Like Virtual Processor Manager 40205, Components 40207 and 40209 of Operating System 40102 multiplex physical Sresources among the virtual resources. As described above, one set of virtual resources may use another set.

One way in which this can happen is indicated by the broken arrows in Figure 402. These arrows show a binding I between Virtual Memory 40206 and Virtual I/O Device 40208. This binding occurs when Virtual Memory 40206 must handle a reference to data contained on a ;-,l'heral I -571device such as a disk drive. To the user of Virtual Memory 40206, all data appears to be available in MEM 10110. In fact, however, the data is stored on peripheral devices such as disk drives, and copied into MEM 10112 when required. When a Process 610 references data that has not been copied into MEM 10112, Virtual Memory 40206 must use IOS 10116 to copy the data into MEM 10112. In order to do this, it uses a Virtual I/O Device S*t 40208 provided by Virtual I/O Manager 40209.

S" 10 b. The--Opelatin-- System--in--S--±1- O 9 For the sake of clarity, Operating System 40102 has been described as though it existed outside of CS 10110.

In fact, however, Operating System 40102 itself uses the resources it controls. In the present embodiment, parts of Operating System 40102 are embodied in JP 10114 hardware devices, parts are embodied in microcode which executes on JP 10114, and parts are embodied in Procedures 602. These Procedures 602 are sometimes S.t called by Processes 610 executing user programs, and sometimes by special Operating System Processes 610 which do nothing but execute operations for Operating System 40102.

The manner in which the components of Operating SSystem 40102 interact may be illustrated by the way in which CS 10110 handles a page fault, a reference to data which is not available in MEM 10110. The first indication that there may be a page fault is an ATU Miss Event Signal. This Event Signal is generated by ATU 10228 in FU 10120 when there is no entry in ATU 10228 for a Logical Descriptor 27116 used in a read or write -572operation. The Event Signal invokes Operating System 40102 microcode, which examines a table in MEM 10112 in order to find whether the data described by Logical Descriptor 27116 has a copy in MEM 10112. If the table indicates that there is no copy, Operating System 40102 microcode communicates the fact of the page fault to an Operating System 40102 Virtual Memory Manager Process 610 and removes Virtual Processor 612 bound to the Process 610 which was executing when the page fault occurred from 10 JP 10114. Some time later, Virtual Memory Manager S. Process 610 is bound to JP 10114. Procedures 602 wl* executed by Virtual Memory Manager Process 610 then initiate the I/O operations required to locate the SA desired data in Secondary Storage 10124 and copy it into MEM 10112. When the data is available in MEM 10112, Operating System 40102 allows Virtual Processor 612 bound to Process 610 which was executing when the page fault occurred to return to JP 10114. Virtual Processor 612 Snrepeats the memory reference which caused the page fault, and since the data is now in MEM 10112, the reference succeeds and execution of Process 610 continues.

c c C. Extended>-Operatnq--5ysem -and--he---en SOpeatin--Svsem--+PH4q-4t In CS 10110, Operating System 40102 is made up of two component operating systems, the Extended Operating System (EOS) and the Kernel Operating System (KOS). The KOS has direct access to the physical resources. It defines a set of primitive virtual resources and multiplexes the physical resources among the primitive virtual resources. The EOS has access to the primitive -573virtual resources defined by KOS, but not to the physical resources. The EOS defines a set of user-level virtual resources and multiplexes the primitive virtual resources defined by KOS among the user level virtual resources.

For example, KOS provides EOS with Processes 610 and Virtual processors 612 and binds Virtual Processors 612 to JP 10114, but EOS decides when a Process 610 is to be created and when a Process 610 is to be bound to a *"rr Virtual Processor 612.

Figure 403 shows the relationship between a user Process 610, EOS, KOS, and the physical resources in CS 10110. Figure 403 shows three levels of interface between executing user Process 610 and JP 10114. The highest level of interface is Procedure Level 40302. At this level, Process 610 interacts with CS 10110 by calling Procedures 602 as specified by the program Process 610 is executing. The calls may be either calls to User Procedures 40306 or calls to EOS Procedures 40307. When Process 610 is executing a Procedure 602, Process 610 produces a stream of SINs. The stream contains two kinds of SINs, S-language SINs 40310 and KOS SINs 40311. Both kinds of SINs interact with CS 10110 at the next level of interface, SIN-level Interface 40309. SINs 40310 and 40311 are interpreted by Microcode 40312 and 40313, and Microinstructions 40315 interact with CS Iw 10110 at the lowest level of interface, JP 10114 j' Interface 40316. As already explained in the discussion of the FU 10120 micromachine, certain conditions in JP 10114 result in Event Signals 40314 which invoke microroutines in S-interpreter Microcode 40312 or KOS 4 I I -574- Microcode 40313. Only Procedure-Level Interface 40302 and SIN-level Interface 40309 are visible to users.

Procedure-level Interface 40309 appears as calls in user Procedures 602 or as statements in user Procedures 602 which compilers translate into calls to EOS Procedures 602. SIN-level Interface 40309 appears as the Name Tables 10335 and SINs in Procedure Objects 608 generated by compilers.

'As Figure 403 indicates, EOS exists only at 10 Procedural Level 40302, while KOS exists at Procedural Level 40302, and SIN Level 40304, and within the microcode beneath SIN Level 40309. The only portion of the operating system that is directly available to user Processes 610 is EOS Procedures 40307. EOS Procedures 40307 may in turn call KOS Procedures 40308. In many cases, an EOS Procedure 40307 will contain nothing more I than the call to a KOS Procedure.40308.

User Procedures 40306, EOS Procedures 40307, and KOS Procedures 40308 all contain S-language SINs 40310. In addition, KOS Procedures 40308 only may contain special KOS SINs 40311. Special KOS SINs 40311 control functions that are not available to EOS Procedures 40307 or User Procedures 40306, and KOS SINs 40311 may therefore not Sappear in Procedures 40306 or 40307. S-language SINs 40310 are interpreted by S-interpreter Microcode 40312, while KOS SINs 40311 are interpreted by KOS Microcode 40313. KOS Microcode 40313 may also be called by Sinterpreter Microcode 40313. Depending on the hardware conditions that cause Event Signals 40314, Signals 40314 -575may cause the execution of either S-interpreter Microcode 40312 or KOS Microcode 40313.

Figure 403 shows the system as it is executing a user Process 610. There are in addition special Processes 610 reserved for KOS and EOS use. These Processes 610 work like user Processes 610, but carry out operating system functions such as process management and virtual memory management. With one exception, EOS Processes 610 call S EOS Procedures 40307 and KOS Procedures 40308, while KOS Processes 610 call only KOS Procedures 40308. The Ct /exception is the beginning of Process 610 execution: KOS performs the KOS-level functions required to begin F'executing a Process 610 and then calls EOS. EOS performs the required EOS level functions and then calls the first User Procedure 40306 in the program Process 610 is executing.

A description of how KOS handles page faults can serve to show how the parts of the system at the JP 2 10114-, SIN-, and Procedure Levels work together. A page I 20 fault occurs when a Process 610 references a data item that has no copy in MEM 10112. The page fault begins as an Event Signal from ATU 10228. The Event Signal invokes H a microroutine in KOS Microcode 40313. If the microroutine confirms that the referenced data item is not in MEM 10112, it records the fact of the page fault in some KOS tables in MEM 10112 and calls another KOS microroutine that unbinds Virtual Processor 612 bound to Process 610 that caused the page fault from JP 10114 and allows another Process 610's Virtual Processor 612 to run. Some time after the page fault, a special operating t i -576system Process 610, the Virtual Memory Manager Process 610, runs and executes KOS Procedures 40309. Virtual Memory Manager Process 610 initiates the I/O operation that reads the data from Secondary Storage 10124 into MEM 10112. When IOS 10116 has finished the operation, Process 610 that caused the page fault can run again and Virtual Memory Manager Proces-s 610 performs an operation which causes Process 610's Virtual Processor 612 to again be bound to JP 10114. When Process 610 resumes Cl,, 10 execution, it again attempts to reference the data. The data is now in MEM 10112 and consequently, the page fault does not recur.

The division of Operating System 40102 into two hierarchically-related operating systems is Is characteristic for CS 10110. Several advantages are gained by such a division: Each of the two operating systems is simpler than a single operating system would be. EOS can concern itself mainly with resource allocation policy and high-level virtual resources, while KOS can concern itself with low-level virtual resources and hardware 41 t control.

Because each operating system is simpler, it is easier to verify that each system's components are performing correctly, and the two systems are therefore more dependable than a single system.

a~ -577- Dividing Operating System 40102 makes it easier to implement different embodiments of CS 10110.

Only the interface provided by EOS is visible to the user, and consequently, the user interface to the system can be changed without altering KOS. In fact, a single CS 10110 may have a number of EOSs, and thereby present different interfaces to different users. Similarly, changes in the hardware affect the 4 4* T: 10 implementation of the KOS, but not the interface that KOS provides EOS. A given EOS can Sr therefore run on more than one embodiment of CS S10110.

A divided operating system is more secure than a single operating system. Physical access to JP 10114 is provided solely by KOS, and consequently, KOS can ensure that users manipulate only those resources to which they have access rights.

S 20 All CSs 10110 will have the virtual resources defined by KOS, while the resources defined by EOS will vary from tto one CS 10110 to another and even within a sngle CS 10110 Consequently, the remainder of the discussion will concern itself with KOS.

The relationship between the KOS and the rest of CS 10110 is governed by four principles: Only the KOS has access to the resources it 4 controls. User calls to EOS may result in EOS S -578calls to KOS, and S-language SINs may result in invocations of KOS microcode routines, but neither EOS nor user programs may directly manipulate resources controlled by KOS.

The KOS is passive. It responds to calls from the EOS, to microcode invocations, and to Event Signals, but it initiates no action on its own.

The KOS is invisible to all system users but the EOS. KOS does not affect the logical behavior of a Process 610 and is noticeable to users only Swith regard to the speed with which a Process S< 610 executes on CS 10110.

As discussed above, KOS manages both physical and virtual resources. The physical resources and some of the virtual resources are visible only within KOS; others of the virtual resources are provided to EOS. Each virtual resource has two main parts: a set of data bases that contain the virtual resource's state, and a set of routines that manipulate the virtual resource. The set l 20 of routines for a virtual resource are termed the resource's manager. The routines may be KOS Procedures 40308, or they may be KOS Microcode 40313. As mentioned, Sin some cases, KOS uses separate Processes 610 to manage Sthe resources.

For the purposes of this specification, the resources managed by KOS fall into two main groups: those associated with objects, and those associated with Processes 610. In the following, first those resources -579associated with objects, and then those associated with Processes 610 are discussed.

B. Obi ects--and-Ob~ ect-Manaeemen-P-t--4t4 The virtual resources termed objects are defined by KOS and manipulated by EOS and KOS. Objects as seen by EOS have five properties: A single UID that identifies the object throughout the object's life and specifies what Logical Allocation Unit (LAU) the object belongs to.

10 A set of attributes that describe the object and r n c t limit access to it.

S Bit-addressable contents. I the present embodiment, the contents may range from 0 to 1 bits in length. Any bit in the contents may be addressed by an offset.

Objects may be created.

Objects may be destroyed.

All of the data and Procedures 602 in a CS 10110 are contained in objects. Any Process 610 executing on a CS 10110 may use a UID-offset address to attempt to access data or Procedures 602 in certain objects on any CS 10110 F, accessible to the CS 10110 on which Process 610 is ;executing. The objects which may be thus accessed by any Process 610 are those having UIDs which are guaranteed unique for all present and future CS 10110. Objects with such unique UIDs thus form a single address space which is at least potentially accessible to any Process 610 executing on any CS 10110. As will be explained in 1 1 -580detail later, whether a Process 610 can in fact access an object in this single address space depends on whether Process 610 has access rights to the object. Other objects, whose UIDs are not unique, may be accessed only by Processes 610 executing on CSs 10110 or groups of CSs 10110 for which the non-unique UID is in fact unique. No two objects accessible to a CS 10110 at a given time may have identical UIDs.

The following discussion of objects will first deal 10 Io with objects as they are seen directly by EOS and et" indirectly by user programs, and then deal with objects as they appear to KOS.

Figure 404 illustrates how objects appear to EOS.

The object has three parts: the UID 40401, the Attributes 40404, and the Contents, 40406. The object's contents reside in a Logical Allocation Unit (LAU), 40405. UID 40401 has two parts: a LAU Identifier (LAUID) 40402 that indicates what LAU 40405 the object is on, and the Object Serial Number (OSN) 40403, which specifies the object in 20 LAU 40405.

The EOS can create an object on a LAU 40405, and given the object's UID 40401, can destroy the object. In addition, EOS can read and change an object's Attributes S40404. Any Process 610 executing on a CS 10110 may reference information in an object by specifying the object's UID 40401 and the bit in the object at which the information begins. At the highest level, addresses in CS 10110 thus consist of a UID 40401 specifying an object -581and an offset specifying the number of bits into the object at which the information begins. As will be explained in detail below, KOS translates such UID-offset addresses into intermediate forms called AON-ofofset addresses for use in JP 10114 and into page numberdisplacement addresses for use in referencing information which has been copied into MEM 10112.

The physical implementation and manipulation of objects is restricted solely to KOS. For instance, tt 10 objects and their attributes are in fact stored in r Secondary Storage 10124. When a program references a Sportion of an object, KOS copies that portion of the object from Secondary Storage 10124 into MEM 10112, and if the portion in MEM 10112 is changed, updates the copy of the object in Secondary Storage 10124. EOS and user programs cannot control the location of an object in Secondary Storage 10124 or the location of the copy of a portion of an object in MEM 10112, and therefore can 1 access the object only by means of KOS.

S 20 While EOS cannot control the physical implementation of an object, it can provide KOS with information that allows KOS to manage objects more effectively. Such information is termed hints. For instance, KOS generally S'i copies a portion of an object into MEM 10112 only if a Process 610 references information in the object.

However, EOS schedules Process 610 execution, and therefore can predict that certain objects will be required in the near future. EOS can pass this information on to KOS, and KOS can use the information to decide what portions of objects to copy into MEM 10112.

f; :II1~L~ I- aC~ -582rrtr r a rrra rt 1 ti a. Obiects--an--UseJ-Proorams-4-Pi-.--45 As stated above, user programs manipulate.objects, but the objects are generally not directly visible to user programs. Instead, user programs use symbols such as variable names or other references to refer to data stored in objects or file names to refer to the objects themselves. The discussion of Namespace has already illustrated how CS 10110 compilers translate variable names appearing in statements in Procedures 602 into 10 Names, indexes of NTEs 30401, how Name Resolve microcode resolves NTE 30401 into Logical Descriptorw 27116, and how ATU 10226 translates Logical Descriptors 27116 into locations in MEM 10112 containing copies of the portions of the objects in which the data represented by the variables resides.

The translation of filenames to UIDs 40401 is accomplished by EOS. EOS maintains a filename translation table which establishes a relationship between a system filename called a Pathname and the UID 40401 of the object containing the file's data, and thereby associates the Pathname with the object. A Pathname is a sequence of ASCII characters which identifies a file to a user of CS 10110. Each pathname in a given CS 10110 must be unique. Figure 405 shows the filename translation table. Referring to that figure, when a user gives Pathname 40501 to the EOS, EOS uses Filename Translation Table 40503 to translate Pathname 40501 into UID 40401 for object 40504 containing the file. An object in CS 10110 may thus be identified in two ways: by means of its UID 40401 or by means of a

I

-583- Pathname 40501. While an object has only a single UID 40401 throughout its life, the object may have many Pathnames 40501. All that is required to change an object's Pathname 40501 is the substitution of one Pathname 40501 for another in the object's Entry 40502 in Filename Translation Table 40503. One consequence of the fact that an object may have different Pathnames 40501 during its life is that when a program uses a Pathname 40501 to identify an object, a user of CS 10110 may make the program process a different object simply by giving the object which formerly had Pathname 40501 which 4* appears in the program a new Pathname 40501 and giving t the next object to be processed the Pathname 40501 which appears in the program.

In the present embodiment, an object may contain only a single file, and consequently, a Pathname 40501 always refers to an entire object. In other embodiments, a Pathname 40501 may refer to a portion of an object, and in such embodiments, Filename Translation Table 40503 20 will associate a Pathname 40501 with a UID-offset address specifying the beginning of the file.

b. u Is-4040-lfFig-a-4of-6 UIDs 40401 way identify objects and other entities in CS 10110. Any entity identified by a UID 40401 has only a single UID throughout its life. Figure 406 is a detailed representation of a CS 10110 UID 40401. UID 40401 is 80 bits long, and has two fields. Field 40402, 32 bits long, is the Logical Allocation Unit Identifier (LAUID). It specifies LAU 40405 containing the object.

LAUID 40402 is further subdivided into two subfields: LAU -584- Group Number (LAUGN) 40607 and LAU Serial Number (LAUSN) 40605, LAUGN 40607 specifies a group of LAUs 40405, and LAUSN 40605 specifies a LAU 40405 in that group.

Purchasers of CS 10110 may obtain LAUGNs 40607 from the manufacturer. The manufacturer guarantees that he will assign LAUGN 40607 given the purchaser to no other CS 10110, and thus these LAUGNs 40607 may be used to form UIDs 40401 which will be unique for all CSs 10110. Field 40604, 48 bits long, is the Object Serial Number (OSN).

It specifies the object in LAU 40405.

UIDs 40401 are generated by KOS Procedures 602.

There are two such Procedures 602, one which generates UIDs 40401 which identify objects, and another which generates UIDs 40401 which identify other entities in CS 10110. The former Procedure 602 is called Generate Object UID, and the latter Generate Non-object UID. The Generate Object UID Procedure 602 is called only by the KOS Create Object Procedure 602. Create Object Procedure S, 602 provides Generate Object UID Procedure 602 with a 20 SI 20 LAUID 40402, and Generate Object UID Procedure 602 returns a UID 40401 for the object. In the present S, embodiment, UID 40401 is formed by taking the current value of the architectural clock, contained in a location Sin MEM 10112, forming an OSN 40403 from the architectural clock's current value, and concatenating OSN 40403 to LAUID 40402.

Generate Non-object UID Procedure 602 may be invoked by EOS to provide a UID 40401 which does not specify an i -585object. Non-object UIDs 40401 may be used in CS 10110 wherever a unique label is required. For example, as will be explained in detail later, all Virtual Processors 612 which are available to CS 10110 have non-object UIDs 40401. All such non-object UIDs 40401 have a single LAUSN 40607, and thus, EOS need only provide a LAUGN 40605 as an argument. Generate Non-object UID Procedure 602 concatenates LAUGN 40605 with the special LAUSN ot 40607, and LAUID 40402 thus produced with an OSN 40403 S 10 obtained from the architectural clock. In other embodiments, OSNs 40403 for both object and non-object UIDs 40401 may be generated by other'means, such as counters.

CS 10110 also has a special UID 40401 called the Null UID 40401. The Null UID 40401 contains nothing but 0 bits, and is used in situations which require a UID value which cannot represent an entity in CS 10110.

c. Object--Attr-butes What a program can do with an object is determined by S 20 the object's Attributes 40404. There are two kinds of Attributes 40404: Object Attributes and Control Attributes. Object Attributes describe the object's r contents;.Control Attributes control access to the

I

f object. Objects may have Attributes 40404 even though they have no Contents 40406, and in some cases, objects may even exist solely for their Attributes 40404.

For the purposes of this discussion, there are two kinds of Object Attributes: the Size Attribute and the Type Attributes.

I j -586- An object's Size Attribute indicates the number of bits that the object currently contains. On each reference to an object's Contents 40406, KOS checks to make sure that the data accessed does not extend beyond the end of the object. If it does, the reference is aborted.

The Type Attributes indicate what kind of information the object contains and how that information may be used. There are three categories of Type Attributes: the Primitive Type Attributes, the Extended Type Attribute, and the Domain of Execution attribute. An object's Primitive Type Attribute indicates whether the object is a data object, a Procedure Object 608, an Extended Type Manager, or an S-interpreter. As their names imply, data objects contain data and Procedure Objects 608 contain Procedures 602. Extended Type Managers (ETMs) are a special type of Procedure Object 608 whose Procedures 608 may perform operations solely on objects called Extended C Type Objects. Extended Type Objects (ETOs) are objects 20 which have an Extended Type Attribute in addition to their Primitive Type Attribute; for details, see the discussion of the Extended Type Attribute below. Sinterpreters are objects that contain interpreters for Slanguages. In the present embodiment, the interpreters consist of dispatch tables and microcode, but in other embodiments, the interpreters may themselves be written in high-level languages. Like the Length Attribute, the Primitive Type Attributes allow KOS to ensure that a S- -587program is using an object correctly. For instance, when the KOS executes a call for a Procedure 602 it checks whether the object specified by the call is a Procedure Object 608. If it is not, the call fails.

d. Attr-ibttes--nd--cces ti.

The remaining Object Attributes and the Control Attributes are all part of CS 10110's Access Control System. The Access Control System is discussed in detail later; here, it is dealt with only to the extent required for the discussion of objects. In CS 10110, an access of an object occurs when a Process 610 fetches SINs contained in a Procedure Object 608, reads data from an object, writes data to an object, or in some cases, when I ,Process 610 transfers control to a Procedure 602. The Access Control System checks whether a Process 610 has the right to perform the access it is attempting. There are two kinds of access in CS 10110, Primitive Access and Extended Access. Primitive Access is access which the Access Control System checks on every reference to an 20 object by a Process 610; Extended Access is access that is checked only on user request. Primitive access checks are performed on every object; extended access checks may ot* be performed only on ETOs, and may be performed only by Procedures 602 contained in ETMs.

The means by which the Access Control System checks a Process 610's access to an object are Process 610's subject and the object's Access Control Lists (ACLs).

L

-588- Each Process 610 has a subject made up of four UIDs 40401. These UIDs 40401 specify the following: The user for whom Process 610 was created. This UID 40401 is termed the principal component of the subject.

Process 610 itself. This UID 40401 is termed the process component.

The domain in which Process 610 is currently executing. This UID 40401 is termed the domain 10 component.

A user-defined subgroup of subjects. This UID 40401 is termed the tag component.

A domain is a group of objects which may potentially be accessed by any Process 610 which is executing a Procedure 602 in one of a group of Procedure Objects 608 or ETMs. Each Procedure Object 608 or ETM has a Domain of Execution (DOE) Attribute. This attribute is a UID 40401, and while a Process 610 is executing a Procedure 602 in that Procedure Object 608 or ETM, the DOE 20 attribute UID 40401 is the domain component in Process 610's subject. The DOE attribute thus defines a group of objects which may be accessed by a Process 610 executing Procedures 602 from Procedure Object 608. The group of objects is called Procedure Object 608's domain. As may 25 be seen from the above definition, a subject's domain component may change on any call to or return from a Procedure 602. The tag component may change whenever the user desires. The principal component and the process component, on the other hand, do not change for the life of Process 610.

r 4 z 7 i ,i ~I~ EC 1 I I, It I.

i it CI I

I.

-589- The ACLs which make up the other half of the Access Control System are attributes of objects. Each ACL consists of a series of Entries (ACLE) and each ACLE has two parts: a Subject Template and a set of Access Privileges. The Subject Template defines a group of subjects, and the set of Access Privileges define the kinds of access that subjects belonging to the group have to the object. To check whether an access to an object is legal, the KOS examines the ACLs. It allows access only if it finds an ACLE whose Subject Template matches the current subject of Process 610 which wishes to make the access and whose set of Access Privileges includes the kind of access desired by Process 610. For example, a Procedure Object 608 may have an ACL with two entries: 15 one whose Subject Template allows any subject access, and whose set of Access Privileges allows only Execute Access, and another whose Subject Template allows only a single subject access and whose set of Access Privileges allows Read, Write, and Execute Access. Such an ACL allows any user of CS 10110 to execute the Procedures 602 in Procedure Object 608, but only a specified Process 610 belonging to a specified user and executing a specified group of Procedures 602 may examine or modify the Procedures 602 in the Procedure Object 608.

There are two kinds of ACLs. All objects have Primitive Access Control Lists (PACLs); ETOs may in addition have Extended Access Control Lists (EACLs). The subject portion of the ACLE is the same in all ACLs; the two kinds of list differ in the kinds of access they control. The access controlled by the PACL is defined by

"U

r -590- KOS and is checked by KOS on every attempt to gain such access; the access controlled by the EACL is defined by the user and is checked only when the user requests KOS to do so.

e. Imp1ementation-of--Obects- 1. Introduction-fti-.-40*7-.--48 The user of a CS 10110 need only concern himself with objects as they have just been described. In order for a Process 610 to reference an object, the object's LAU 40405 must be accessible from CS 10110 upon which Process 610 is running, Process 610 must know the object's UID r" 40401, and Process 610's currernt subject must have the right to access the object in the desired manner. Process 610 need know neither how the object's Contents 40406 and 15 Attributes 40404 are stored on CS 10110's physical devices nor the methods CS 10110 uses to make the object's Contents 40406 and Attributes 40404 available to Process 610.

S, ,The KOS, on the other hand, must implement objects on the physical devices that make.up CS 10110. In so doing, ste it must take into account two sets of physical Slimitations: t t In logical terms, all CSs 10110 have a single logical memory, but the physical implementation t1 25 of memory in the system is hierarchical: a given SCS 10110 has rapid access to a relatively small MEM 10112, much slower access to a relatively large amount of slow Secondary Storage 10124, and very slow access to LAUs 40405 on other accessible CSs 10110.

-591- UIDs 40401, and even more, subjects, are too large to be handled efficiently on JP 10114's internal data paths and in JP 10114's registers.

The means by which the KOS overcomes these physical limitations will vary from embodiment to embodiment.

Here, there are presented first an overview and then a detailed discussion of the means used in the present embodiment.

The physical limitations of the memory are overcome by means of a Virtual Memory system. The Virtual Memory System creates a one-level logical memory by automatically bringing copies of those portions of (objects required by executing Processes 610 into MEM 10112 and automatically copying altered portions of objects from MEM 10112 back to Secondary Storage 10124.

Objects thus reside primarily in Secondary Storage 10124, but copies of portions of them are made available in MEM 10112 when a Process 610 makes a reference to them.

Besides bringing portions of objects into MEM 10112, when j t 20 required, the Virtual Memory System keeps track of where in MEM 10112 the portions are located, and when a Process 610 references a portion of an object that is in MEM 10112, the Virtual Memory System translates the reference into a physical location in MEM 10112.

JP 10114's need for smaller object identifiers and S' subject identifiers is satisfied by the use of internal identifiers called Active Object Numbers (AONs) and Active Subject Numbers (ASNs) inside JP 10114. Each time a UID 40401 is moved from MEM 10112 into JP 10114's ,4 -592registers, it is translated into an AON, and the reverse translation takes place each time an AON is moved from a JP 10114's registers to MEM 10112. Similarly, the current subjects of Processes 610 which are bound to Virtual Processors 612 are translated from four UIDs 40401 into small integer ASNs, and when Virtual Processor 612 is bound to JP 10114, the ASN for the subject belonging to Virtual Processor 612's Process 610 is placed in a JP 10114 register. The translations from UID 40401 to AON and vice-versa, and from subject to ASN are performed by KOS.

When KOS translates UIDs 40401 to AONs and viceversa, it uses AOT 10712. An AOT 10712 Entry (AOTE) for an object contains the object's UID 40401, and the AOTE's index in AOT 10712 is that object's AON. Thus, given an object's AON, KOS can use AOT 10712 to determine the object's UID 40401, and given an obect's UID 40401, KOS can use AOT 10712 to determine the object's AON. If the object has not been referenced recently, there may be no S 20 AOTE for the object, and thus no AON for the object's UID I.i 40401. Objects that have no AONs are called inactive objects. If an attempt to convert a UID 40401 to an AON t' reveals that the object is inactive, an Inactive Object Fault results and KOS must activate the object, that is, it must assign the object an AON and make an AOTE for it.

KOS uses AST 10914 to translate subjects into ASN's.

When a Process 610's subject changes, AST 10914 provides Process 610 with the new subject's ASN. A subject may 4A, -593presently have no ASN associated with it. Such subjects are termed inactive subjects. If a subject is inactive, an attempt to translate the subject to an ASN causes KOS to activate the subject, that is, to assign the subject an ASN and make an entry for the subject in AST 10914.

In order to achieve efficient execution of programs by Processes 610, KOS accelerates information that is frequently used by executing Processes 610. There are two stages of acceleration: Tables that contain the information are wired into MEM 10112, that is, the Virtual Memory System never uses MEM 10112 space reserved for the tables for other purposes.

Special hardware devices in JP 10114 contain portions of the information in the tables.

MHT 10716, AOT 10712, and AST 10914 are examples of the first stage of acceleration. As previously mentioned, these tables are always present in MEM 10112. Address Translation Unit (ATU) 10228 is an example of the second 20 stage. As previously explained, ATU 10228 is a hardware cache that contains copies of the most recently used MHT S.4 10716 entries. Like MHT 10716, it translates AON offset r *addresses into the MEM 10112 locations that contain copies of the data that the UID-offset address e: 25 corresponding to the AON-offset address refers to. ATU 10228 is maintained by KOS Logical Address Translation (LAT) microcode.

Figure 407 shows the relationship between ATU 10228, MEM 10112, MHT 10716, and KOS LAT microcode 40704. When JP 10114 makes a memory reference, it passes AON-offset -594- Address 40705 to ATU 10228. If ATU 10228 contains a copy of MHT 10716's entry for Address 40705, it immediately produces the corresponding MEM 10112 Address 40706 and transmits the address to MEM 10112. If there is no copy, ATU 10228 produces an ATU Miss Event Signal which invokes LAT microcode 40704 in JP 10114. LAT microcode 40704 obtains the MHT entry that corresponds to the AON-offset address from MHT 10716, places the entry in ATU 10228, and returns. JP 10114 then repeats the reference. This o0 time, there is an entry for the reference, and ATU 10228 translates the AON address into the address of the copy of the data contained in MEM 10112.

t The relationship between KOS table, hardware cache, and microcode just described is typical for the present embodiment of CS 10110. The table (in this case, MHT 10716) is the primary source of information and is maintained by the Virtual Memory Manager Process, while the cache accelerates portions of the table and is n maintained by KOS microcode that is invoked by event 20 signals from the cache.

t AOT 10712, AST 10914, and MHT 10716 share another 4, characteristic that is typical of the present embodiment S" of CS 10110: the tables are constructed in such a fashion that the table entry that performs the desired |1 25 translation is located by means of a hash function and a hash table. The hash function translates the large UID 40401, subject, or AON into a small integer. This integer is the index of an entry in the hash table. The -595contents of the hash table entry is an index into AOT 10712, AST 10914, or MHT 10716, as the case may be, and these tables are maintained in such a fashion that the entry corresponding to the index provided by the hash table is either the entry that can perform the desired translation or contains information that allows KOS to find the desired entry. The entries in the tables furthermore contain the values they translate.

Consequently, KOS can hash the value, find the entry, and then check whether the entry is the one for the hashed value. If it is not, KOS can quickly go from the entry or located by the hash table to the correct entry.

Figure 408 shows how hashing works in AST 10914 in the present embodiment. In the present embodiment, Subject 40801, the principal, process, and domain components of the current subject, are input into Hash Function 40802. Hash Function 40802 produces the index of an entry in ASTHT 10710. ASTHT Entry 40504 in turn o contains the index of an Entry (ASTE) 40806 in AST 10914 These ASTE 40806 indexes are ASNs. ASTE 40806 contains the principal, process, and domain components of some subject and a link field pointing to ASTE 40806'. ASTE 40806' has 0 in its link field, which indicates that it is the last link in the chain of ASTES begining with ASTE S 25 40806. If the hashing of a subject yields ASTE 40806, KOS compares the subject in ASTE 40806 with the hashed subject; if they are identical, ASTE 40806's index in AST 10914 is the subject's ASN. If they are not identical, is the subject's ASN; otherwise, ASTE 40806' is the last entry in the chain, and consequently, there is no ASTE 40806 and no ASN for the hashed subject.

In the following, we will discuss the implementation of objects in the present embodiment in detail, beginning with the implementation of objects in Secondary Storage 10124 and proceeding then to CS 10110's Active Object Management System, the Access Control System, and the Virtual Memory System.

2. Oble.ct.s: in-Secondarty-Sto-race--l--24 4.09i-410-) As described above, objects are collected into LAUs 40405. The objects belonging to a LAU 40405 are stored in Secondary Storage 10124. Each LAU 40405 contains an object whose contents are a table called the Logical Allocation Unit Directory (LAUD). As its name implies, S, 20 the LAUD is a directory of the objects'in LAU 40405.

x| Each object in LAU 40405, including the object containing the LAUD, has an entry in the LAUD. Figure 409 shows the I Y relationship between Secondary Storage 10124, LAU 40405, the LAUD, and objects. LAU 40405 resides on a number of Storage Devices 40904. LAUD Object 40902' in LAU 40405 contains LAUD 40903. Two LAUDEs 40906 are shown. One contains the attributes of LAUD Object 40902 and the location of its contents, and the other contains the attributes of LAUD Object 40902' containing LAUD 40903 a -597and the location of its contents.

KOS uses a table called the Active LAU Table (ALAUT) to locate the LAUD belonging to LAU 40405. Figure 410 illustrates the relationship between ALAUT 41001, ALAUT Entries 41002, LAUs 40405, and LAUD Objects 40902'. Each LAU 40405 accessible to CS 10110 has an Entry (ALAUTE) 41002 in ALAUT 41001. ALAUTE 41002 for LAU 40405 includes LAU 40405's LAUID 40402 and UID 40401 of LAU 40705's LAUD Object 40902'. Hence, given an object's UID 40401, KOS can use UID 40401's LAUID 40402 to locate ALAUTE 41002 for the object's LAU 40405, and can use ALAUTE 41002 to locate LAU 40405's LAUD 40903. Once LAUD 40903 has been found, OSN portion 40402 of the object's UID 40401 provides the proper LAUDE 40906, and LAUDE 40906 contains object's attributes and the location of its contents.

LAUD 40903 and the Procedures 602 that manipulate it belong to a part of FOS termed the Inactive Object Manager. The following discussion of the Inactive Object Manager will begin with the manner in which an object's 20 contents are represented on Secondary Storage 10124, will then discuss LAUD 40903 in detail, and conclude by discussing the operations performed by Inactive Object Manager Procedures 602.

a.a. Representation--of--an--Obect-LS Contents-onSeconda y-Storaqe In general, the manner in which an object's contents are represented on Secondary Storage 10124 depends completely on the Secondary Storage 10124. If a LAU 40405 is made up of disks, then the object's contents r t -598will be stored in disk blocks. As long as KOS can locate the object's contents, it makes no difference whether the storage is contiguous or non-contiguous.

In the present embodiment, the objects' contents are stored in files created by the Data General Advanced Operating System (AOS) procedures executing on IOS 10116. These procedures manage files that contain objects' contents for KOS. In future CSs 10110, the representation of an object's contents on Secondary Storage 10124 will be managed by a portion of KOS.

b.b. LAUD--4 90----ff Figure 411 is a conceptual illustration of LAUD 40903. LAUD 40903 has three parts: LAUD Header 41102, t Master Directory 41105, and LAUD Entries (LAUDEs) 40906.

LAUD Header 41102 and Master Directory 41105 occupy fixed locations in LAUD 40903, and can therefore always be located from the UID 40401 of LAUD 40903 given in ALAUT 41001. The locations of LAUDEs 40906 are not fixed, but the entry for an individual object can be located from r' 20 Master Directory 41105.

Turning first to LAUD Header 41102, LAUD Header 41102 contains LAUID 40402 belonging to LAU 40405 to which LAUD 40903 belongs and OSN 40403 of LAUD 40903. As will be explained in greater detail below, KOS can use OSN 40403 to find LAUDE 40906 for LAUD 40903.

Turning now to Master Directory 41105, Master Directory 41105 translates an object's OSN 40403 into the location of the object's LAUDE 40906. Master Directory 41105 contains one Entry 41108 for each object in LAU 40505. Each Entry has two fields: OSN Field 41106 and -599- Offset Field 41107. OSN Field 41106 contains OSN 40403 for the object to which Entry 41108 belongs; Offset Field 41107 contains the offset of the object's LAUDE 40906 in LAUD 40903. KOS orders Entries 41108 by increasing OSN 40403, and can therefore use binary search means to find Entry 41108 containing a given OSN 40403. Once Entry 41108 has been located, Entry 41108's Offset Field 41107, combined with LAUD 40903's OSN 40403, yields the UIDoffset address of the object's LAUDE 40906.

Once KOS knows the location of LAUDE 40906 it can determine an object's Attributes 40404 and the location of its Contents 40406. Figure 411 gives only an overview of LAUDE 40906's general structure. LAUDE 40906 has three components: a group of fields of fixed size 41109 that are present in every LAUDE 40906, and two variablesized components, one, 41139, containing entries belonging to the object's PACL, and another, 41141, containing the object's EACL.

As the preceding descriptions of the LAUD's S 20 components imply, the number of LAUDEs 40906 and Master Direcory Entries 41108 varies with the number of objects in LAU 40405. Furthermore, the amount of space required t for an object's EACL and PACL varies from object to object. KOS deals with this problem by including Free Space 41123 in each LAUD 40903. When an object is created, or when an object's ACLs are expanded, the Inactive Object Manager expands LAUD 40903 only if there is no available Free Space 41123; if there is Free Space i).

-600- 41123, the Inactive Object Manager takes the necessary space from Free Space 41123; when an object is deleted or an object's ACLs shortened, the Inactive Object Manager returns the unneeded space to Free Space 41123.

Figure 412 is a detailed representation of a single LAUDE 40906. Figure 412 presents those fields of LAUDE 40906 which are common to all embodiments of CS 10110; fields which may vary from embodiment to embodiment are ignored. Starting at the top of Figure 412, Structure Version Field 41209 contains information by which KOS can determine which version of LAUDE 40906 it is dealing c with. Size Field 41211 contains the Size Attribute of the object to which LAUDE 40906 belongs. The Size Attribute specifies the number of bits currently contained in the object. Lock Field 41213 is a KOS lock. As will be explained in detail in the discussion of Processes 610, Lock Field 41213 allows only one Process 610 to read or write LAUDE 40906 at a time, and s therefore keeps one Process 610 from altering LAUDE 40906 while another Process 610 is reading LAUDE 40906. File Identifier 41215 contains a system identifier for the file which contains the Contents 40406 of the object to which LAUDE 40906 belongs. The form of File Identifier t l t s 41215 may vary from embodiment to embodiment; in the 25 present embodiment, it is an AOS system file identifier.

UID Field 41217 contains UID 40401 belonging to LAUDE 1C 40906's object. Primitive Type Field 41219 contains a i value which specifies the object's Primitive Type. The L I~i: rr* -601object may be a data object, a Procedure Object 608, an ETM, or an S-interpreter object. AON Field 41221 contains a valid value only when LAUDE 40906's object is active, has an entry in AOT 10712. AON Field 41221 then contains the object's AON. If the object is an ETO, Extended Type Attribute Field 41223 contains the UID 40401 of the ETO's ETM. Otherwise, it contains a Null UIU 40401. Similarly, if the object is a Procedure Object 608 or an ETM, Domain of Execution Attribute Field 41225 contains the object's Domain of Execution Attribute.

The remaining parts of LAUDE 40906 belong to the Acces Control System and will be explained in detail in Sthat discussion. Attribute Version Number Field 41227 contains a value indicating which version of ACLEs this LAUDE 40906 contains, PACL Size Field 41229 and EACL Size Field 41231 contain the sizes of the respective ACLs, PACL Offset Field 41233 and EACL Offset Field 41235 contain the offsets in LAUD 40903 of additional PACLEs 41139 and EACLEs 41141, and fixed PACLEs 41237 contains the portion of the PACL which is always included in LAUDE 40906.

3. Active- Objects-F-Pi--4-93 An active object is an object whose UID 40401 has an AON associated with it. In the present embodiment, each 1 25 CS 10110 has a set of AONs. KOS associates these AONs with UIDs 40401 in such fashion that at any given moment, an AON in a CS 10110 represents a single UID 40401.

Inside FU 10120, AONs are used to represent UIDs CS -602- 10110. In the present embodiment, the AON is represented by 14 bits. A 112-bit UID-offset address (80 bits for UID 40401 and 32 for the offset) is thus represented inside FU 10120 by a 46-bit AON-offset address (14 bits for the AON and 32 bits for the offset).

A CS 10110 has far fewer AONs than there are UIDs 40401. KOS multiplexes a CS 10110's AONs among those objects that are being referenced by CS 10110 and therefore require AONs as well as UIDs 40401. While a given AON represents only a single UID 40401 at any given time, at different times, a UID 40401 may have different AONs associated with it.

i Figure 413 provides a conceptual representation of the relationship between AONs and UIDs 40401. Each CS 10110 has potential access to 2**80 UIDs 40401. Some of these UIDs, however, represent entities other than objects, and others are never associated with any entity. Each CS 10110 also has a set of AONs 41303 available to it. In the present embodiment, this set may have up to 2**14 values. Since the AONS are only used S internally, each CS 10110 may have the same set of AONs 41303. Any AON 41304 in set of AONs 41303 may be associated with a single UID 40401 in set of object UIDs :41301. At different times, an AON 41304 may be 25 associated with different UIDs 40401.

4 As mentioned above, KOS associates AONs 41304 with UIDs 40401. It does so by means of AOT 10712. Each AOT entry (AOTE) 41306 in AOT 10712 associates a UID 40401 -603with an AON 41304. AON 41304 is the index of AOTE 41306 which contains UID 40401. Until AOTE 41306 is changed, the AON 41304 which is the index of AOTE 41306 containing UID 40401 represents UID 40401. AOT 10712 also allows UIDs 40401 to be translated into AONs 41303 and viceversa. Figure 413 illustrates the process for UID-offset Address 41308 and AON-offset Address 41309. AOTE 41306 associates AON 41304 in AON-offset Address 41309 with UID 40401 in UID-offset Address 41308, and Addresses 41308 10 and 41309 have the same Offset 41307. Consequently, AONoffset Address 41309 represents UID-offset Address 41308 1 inside JP 10114. Since both addresses use the same Offset, Address 41309 can be translated into address 41308 by translating Address 41309's AON 41304 into Address 41308's UID 40401, and Address 41308 can be translated into Address 41309 by the reverse process. In both cases, the translation is performed by finding the t proper AOTE 41306.

The process by which an object becomes active is called object activation. A UID-offset Address 41308 cannot be translated into an AON-offset Address 41309 unless the object to which UID 40401 of UID-offset Address 41308 belongs is active. If a Process 610 o' attempts to perform such a translation using a UID 40401 25 belonging to an inactive object, an Inactive Object Fault occurs. KOS handles the fault by removing Process 610 that attempted the translation from JP 10114 until a special KOS Process called the Object Manager Process has activated the object. After the object has been -604activated, Process 610 may return to JP 10114 and complete the UID 40401 to AON 41304 translation.

The portion of KOS that manages active objects is called the Active Object Manager (AOM). Parts of the AOM are Procedures 602 and parts of it are microcode routines. The high-level language components of the AOM may be invoked only by KOS Processes 610. KOS Active Object Manager Process 610 performs most of the functions involved in active object management.

a.a. UID- 40401 to--AON-41-04 SGenerally speaking, in CS 10110, addresses stored in I t MEM 10112 and Secondary Memory 10124 are stored as UIDoffset addresses. The only form of address that FU 10120 can translate into a location in MEM 10112 is the AONoffset form. Consequently, each time an address is loaded from MEM 10112 into a FU 10120 register, the address must Sbe translated from a UID-offset address to an AON-offset address. The reverse translation must be performed each time an address is moved from a FU 10120 register back into memory.

Such translations may occur at any time. For example, a running Virtual Processor 612 performs such a translation when the Process 610 being executed by Virtual Processor 612 carries out an indirect memory reference. An indirect memory reference is a reference which first fetches a pointer, that is, a data item Whose value is the address of another data item, and then uses the address contained in the pointer to fetch the data itself. In CS 10110, pointers represent UID-offset 4 b hL -605addresses. Virtual Processor 612 performs the indirect memory reference by fetching the pointer from MEM 10112, placing it in FU 10120 registers, translating UID 40401 represented by the pointer into AON 41304 associated with it, and using the resulting AON-offset address to access the data at the location specified by the address.

Most such translations, however, occur when Virtual Processor 612 state is saved or restored. For instance, when one Process 610's Virtual Processor 612 is removed t o 10 from JP 10114 and another Process 610's Virtual Processor 612 is bound to JP 10114, the state of Virtual Processor i 612 being removed from JP 10114 is stored in memory, and Sthe state of Virtual Processor 612 being bound to JP 10114 is moved into JP 10114's registers. Because only UID-offset addresses may be stored in memory, all of the AON-offset addresses in the state of Virtual Processor 612 which is being removed from JP 10114 must be translated into UID-offset addresses. Similarly, all of the UID-offset addresses in the state of Virtual Processor 612 being bound to JP 10114 must be translated 'into AON-offset addresses before they can be loaded into FU 10120 registers.

C. The *Access-eont-ro-l--Sytem \As mentioned in the introduction to objects, ach time 25 a Process 610 accesses data or SINs in an object$ the KOS Access Control System checks whether Process 610's current subject has the right to perform the kind of access that Process 610 is attempting. If Process 610's -606current subject does not have the proper access, the Access Control System aborts the memory operation which Process 610 was attempting to carry out. The following discussion presents details of the implementation of the Access Control System, beginning with subjects, then proceeding to subject templates, and finally to the means used by KOS to accelerate access checking.

a. Subiects A Process 610's subject is part of Process 610's state and is contained along with other state belonging site to Process 610 in an object called a Process Object.

Process Objects are dealt with at length in the detailed jo* *discussion of Processes 610 which follows the discussion of objects. While a subject has, as mentioned above, four components, the principal component, the process component, the domain component, and the tag component, the Access Control System in the present embodiment of CS 10110 assigns values to only the first three components and ignores the tag component when checking access.

In the present embodiment, the UIDs 40401 which make up the components of a Process 610's subject are the UIDs 40401 of objects containing information about the entities represented by the UIDs 40401. The principal t 4 *component's UID 40401 represents an object called the 25 Principal Object, The Principal Object contains information about the user for whom Process 610 was created. For example, the information might concern what acass rights the user had to the resources of CS 10110, -607or it might contain records of his use of CS 10110. The process component's UID 40401 represents the Process Object, while the domain component's UID 40401 represents an object called the Domain Object. The Domain Object contains information which must be accessible to,any Process 610 whose subject has the Domain Object's UID 40401 as its domain component. Other embodiments of CS 10110 will use the tag component of the subject. In these embodiments, the tag component's UID 40401 is the UID 40401 of a Tag Object containing at least such information as a list of the subjects which make up the t group of subjects represented by the tag component's UID.

o b. Domains As stated above, the subject's domain component is the domain of execution attribute belonging to the Procedure Object 608 or ETM whose code is being executed when the access request is made. The domain component of Se l the subject thus gives Process 610 to which the subject 4, belongs potential access to the group of objects whose ACLs have ACLEs with subject templates containing domain components that match the DOE attribute. This group of objects is the domain defined by the Procedure Object 608 or ETM's DOE attribute. When a Process 610 executes a *t "Procedure 602 from a Procedure Object 608 or ETM with a 2 25 given DOE attribute, Process 610 is said to be executing in the domain defined by that DOE attribute. As may be inferred from the above, different Procedure Objects 608 or ETMs may have the same DOE attribute, and objects may \41 -608have ACLEs which make them members of many different domains.

In establishing a relationship between a group of Procedure Objects 608 and another group of objects, a domain allows a programmer using CS 10110 to ensure that a given object is read, executedp or modified only by a certain set of Procedures 602. Domains may thus be used to construct protected subsystems in CS 10110. One example of such a protected subsystem is KOS itself: the objects in CS 10110 which contain KOS tables all have ACLs whose domain template components match only the DOE which represents the KOS domain. The only Procedure Objects 608 and ETMs which have this DOE are those which contain KOS Procedures 602, and consequently, only KOS Procedures 602 may manipulate KOS tables.

Since an object may belong to more than one domain, a programmer may use domains to establish hierarchies of access. For example, if some of the objects in a first domain belong both to the first domain and a second domain, and the second domain's objects all also belong to the first domain, then Procedures 602 contained in Procedure Objects 608 ehome DOEs define the first domain may access any object in the first domain, including those which also belong to the second domain, while those S 25 from Procedure Objects 608 whose DOEs define the second domain may access only those objects in the second domain.

co-i K -O re e onl KOS -609- C. Access--Contro-l---st As previously mentioned, the Access Control System compares the subject belonging to Process 610 making an access to an object and the kind of access Process 610 desires to make with the object's ACLs to determine whether the access is legal. The following discussion of the ACLs will first deal with Subject Templates, since they are common to all ACLs, and then with PACLs and EACLs.

1. Sub ect--Temp-laes--P'g-.-46 Figure 416 shows Subject Templates, PACL Entries t* (PACLEs), and EACL Entries (EACLEs). Turning first to ,i o" the Subject Templates, Subject Template 41601 consists of four components, Principal Template 41606, Process Template 41607, Domain Template 41609, and Tag Template 41611. Each template has two fields, Flavor Field 41603, and UID Field 41605. Flavor Field 41603 indicates the way in which the template to which it belongs is to match the corresponding component of the subject for Process 610 attempting the access. Flavor Field 41603 may have one of three values: match any, match one, match group.

If Flavor Field 41603 has the value match any, any subject component UID 40401 matches the template, and the S Access Control System does not examine UID Field 41605.

#14* 25 If Flavor Field 41603 has the value match one, then the corresponding subject component must have the same UID 40401 as the one contained in UID Field 41605. If Flavor Field 41603 has the value match group, finally, then UID S k J ~t" -610- Field 41605 contains a UID 40401 of an object containing information about the group of subject components which the given subject component may match.

2. Prim-it ve--Access--Cont-ro-L-s-s-+PAeei PACLs are made up of PACLEs 41613 as illustrated in Figure 416. Each PACLE 41613 has two parts: a subject template 41601 and an Access Mode Bits Field 41615. The values in Access Mode Bits Field 41615 define 11 kinds of access. The eleven kinds fall into two groups: Primitive Data Access and Primitive Non-data Access. Primitive Data Access controls what the subject may do with the object's Contents 40406; Primitive Non-data Access controls what the subject may do with the object's Attributes 40404.

St, There are three kinds of Primitive Data Access: Read Access, Write Access, and Execute Access. If a subject f has Read Access, it can examine the data contained in the tL object; if the subject has Write Access, it can alter the data contained in the object; if it has Execute Access, I it can treat the data in the object as a Procedure 602 and attempt to execute it. A subject may have none of these kinds of access, or any combination of the kinds.

On every reference to an object, the KOS checks whether the subject performing the reference has the required Primitive Data Access.

Primitive Non-data Access to an object is required «only to set or read an object's Attributes 40404, and is checked only when these operations are performed. The

L

-611kinds of Non-data Access correspond to the kinds of Attributes 40404: Attributes Kind of Access Object Attributes Primitive Control Attributes Extended Control Attributes get object attributes set object attributes get primitive control attributes set primitive control attributes get extended control attributes set extended control attributes 444r 4 4 c LC C Ce 4 44 44 1 4444r 1 44 44 .4 4 *r 4; 4 4 44 444* 4 #4 44 4 ETM Access use as ETM create ETO 20 The access rights for object attributes allow a subject to get and set the object attributes described previously. The access rights for primitive and extended control attributes allow a subject to get and set an object's PACL and EACL respectively.

25 An object may have any number of PACLEs 41613 in its PACL. The first five PACLEs 41613 in an object's PACL are contained in fixed PACLE Field 41237 of LAUDE 40906 for the object; the remainder are stored in LAUD 40903 at the 7 "--r"lc r sr

C

r I: I I II -612location specified in PACL Offset Field 41233 of LAUDE 40906.

3. APAM---109+8--and--Protecti on-eache 10234---(Fig--42-1 Primitive non-data access rights are checked only when users invoke KOS routines that require such access rights, and extended access rights are checked only when users request such checks. Primitive data access rights, on the other hand, are checked every time a Virtual Processor 612 makes a memory reference while executing a Process 610. The KOS implementation of primitive data access right checking therefore emphasizes speed and efficiency. There are two parts to the implementation: APAM 10918 in MEM 10112, and Protection Cache 10234 in JP 15 10114. APAM 10918 is in a location in MEM 10112 known to KOS microcode. APAM 10918 contains primitive data access information copied from PACLEs 41613 which belong to active objects and whose Subject' Template 41601 matches an active subject. Protection Cache 10234, in turn, contain copies of the information in APAM 10918 for the active subject of Process 610 whose Virtual Processor 612 is currently bound to JP 10114 and active objects referenced by Process 610. A primitive data access check in CS 10110 begins with Protection Cache 10234, and if the information is not contained in Protection Cache 10234, proceeds to APAM 10918, and if it is not there, finally, to the object's PACL. The discussion which follows begins with APAM 10918.

.111 >4 j" r 1 -613- Figure 421 shows APAM 10918. APAM 10918 is organized as a two-dimensional array. The array's row indexes are AONs 41304, and its column indexes are ASNs. There is a row for each AON 41304 in CS 10110, and a column for each ASN. In Figure 421, only a single'row and column are shown. Any primitive data access linformation in APAM 10918 for the object represented by AON 41304 j is contained in Row 42104, while Column 42105 contains any primitive data access information in APAM 10918 for the subject represented by ASN k. APAM Entry (APAME) 42106 is at the intersection of Row 42104 and Column 42105, and t thus contains the primitive data access information from t that PACLE 41613 belonging to the object represented by j AON 41304 j whose Subject Template 41601 matches the subject represented by ASN k.

An expanded view of APAME 42106 is presented beneath the epresentation of APAM 10918. APAME 42106 contains four 1-bit fields. The bits represent the kinds of primitive data access that the subject represented by 20 APAME 42106's column index has -to the object represented by APAME 42106's row index.

S 4. Field 42107 is the Valid Bit. If the Valid Bit is set, APAME 42106 contains whatever primitive data access information is available for the 25 subject represented by the column and the object Srepresented by the row. The remaining fields in APAME 42106 are meaningful only if Valid Bit 42107 is set.

i K r ,1 -614- Field 42109 is the Execute Bit. If it is set, APAME 42106's subject has Execute Access to APAME 42106's object.

Field 42111 is the Read Bit. If it is set, APAME 42106's subject has Read Access to APAME 42106's object.

Field 42113 is the Write Bit. If it is set, APAME 42106's subject has Write Access to APAME 42106's object.

Any combination of bits in Fields 42109 through 42113 may be set. If all of these fields are set to 0, APAME 42106 rindicates that the subject it represents has no access to the object it represents.

KOS sets APAME 42106 for an ASN and an AON 41304 the first time the subject represented by the ASN references the object represented by AON 41304. Until APAME 42106 is set, Valid Bit 42107 is set to 0. When APAME 42106 is set, Valid Bit 42107 is set to 1 and Fields 42109 through 1. 42113 are set according to the primitive data access 20 information in the object's PACLE 41613 whose Subject Template 41601 matches the subject. When an object is deactivated, Valid Bits 42107 in all APAMEs 42106 in the row belonging to the object's AON 41304 are set to 0; similarly, when a subject is deactivated, Valid Bits 42107 in all APAMEs 42106 in the column belonging to the subject's ASN are set to 0.

s -615- 4. Protection--Cache--i0-234--and Prtection- Check4in-- -fi 2-2 The final stage in the acceleration of protection information is Protection Cache 10234 in JP 10114. The details of the way in which Protection Cache 10234 functions are presented in the discussion of the hardware; here, there are discussed the manner in which Protection Cache 10234 performs access checks, the relationship between Protection Cache 10234, APAM 10918, and AOT 10712, and the manner in which KOS protection cache microcode maintains Protection Cache 10234.

Figure 422 is a block diagram of Protection Cache 10234, AOTE 10712, APAM 10918, and KOS Microcode 42207 i which maintains Protection Cache 10234. Each time JP 15 10114 makes a memory reference using a Logical Descriptor S.i* 27116, it simultaneously presents Logical Descriptor j27116 and a Signal 42208 indicating the kind of memory operation to Protection Cache 10234 and ATU 10228.

S. Entries 42215 in Protection Cache 10234 contain primitive 4.90. 20 data access and length information for objects previously referenced by the current subject of Process 610 whose Virtual Processor 612 is currently bound to JP 10114. On every memory reference, Protection Cache 10234 emits a Valid/invalid Signal 42205 to MEM 10112. If Protection Cache 10234 contains no Entry 42215 for AON 41304 i contained in Logical Descriptor 27116's AON field 27111, if Entry 42215 indicates that the subject does not have the type of access required by Process 610, or if the sum -616of Logical Descriptor 27116's OFF field 27113 and LEN field 27115 exceed the object's current site, Protection Cache 10234 emits an Invalid Signal 42205. This signal causes MEM 10112 to abort the memory reference.

Otherwise, Protection Cache 10234 emits a Valid Signal 42205 and MEM 10112 executes the memory reference.

When Protection Cache 10234 emits an Invalid Signal 42205, it latches Logical Descriptor 27116 used to make the reference into Descriptor Trap 20256, the memory command into Command Trap 27018, and if it was a write operation, the data into Data Trap 20258, and at the same time emits one of two Event Signals to KOS microcode.

Illegal Access Event Signal 42208 occurs when Process 610 making the reference does not have the proper access rights or the data referenced extends beyond the end of the object. Illegal Access Event Signal 42208 invokes KOS microcode 42215 which performs a Microcode to Software Call 42217 (described in the discussion of Calls) to KOS Access Control System Procedures 602 and 20 passes the contents of Descriptor Trap 20256, Command Trap 27018, the ASN of Process 610 (contained in a register MGR's 10360), and if necessary, Data Trap 20258 to these Procedures 602. Thse Procedures 602 inform EOS of the protection violatior, and EOS can then remedy it.

Cache Miss Event Signal 42206 occurs when there is no 54 Entry 42215 for AON 41304 in Protection Cache 10234.

Cache Miss Event Signal 42206 invokes KOS Protection Cache Miss Microcode 42207, which constructs missing -617- Protection Cache Entry 42215 from information obtained from AOT 10712 and APAM 10918. If APAM 10918 contains no entry for the current subject's ASN and the AON of the object being referenced, Protection Cache Miss Microcode 42207 performs a Microcode-to-software Call to KOS Access Control System Procedures 602 which go to LAUDE 40906 for the object and copy the required primitive data access information from the PACLE 41613 belonging to the object whose Subject Template 41601 matches the subject attempting the reference into APAM 10918. The KOS Access Control System Procedures 602 then return to Cache Miss r t Microcode 42207, which itself returns. Since Cache Miss 'b Microcode 41107 was invoked by an Event Signal, the return causes JP 10114 to reexecute the memory reference which caused the protection cache miss. If Protection Cache 10234 was loaded as a result of the last protection t r cache miss, the miss does not recur; if Protection Cache 10234 was not loaded because the-required information was not in APAM 10918, the miss recurs, but since the 20 information was placed in APAM 10918 as a result of the 4 S* previous miss, Cache Miss Microcode 42207 can now construct an Entry 42215 in Protection Cache 10234. When Cache Miss Microcode 42207 returns, the memory reference is again attempted, but this time Protection Cache 10234 25 contains the information and the miss does not recur.

y 1 jCache Miss Microcode 42207 creates a new Protection Cache Entry 42215 and loads it into Protection Cache 10234 as follows: Using AON 41304 from Logical Descriptor 27116 latched into Descriptor Trap 20255 wheri the memory reference which caused the miss was executed and the -618current subject's ASN, contained in GR's 10360, Cache Miss Microcode locates APAME 42106 for the subject represented by the ASN and the object represented by AON 41304 and copies the contents of APAME 42106 into a JP 10114 register which may serve as a source for JPD Bus 10142. It also uses AON 41304 to locate AOTE 41306 for I the object and copies the contents of Size Field 41519 into another JP 10114 register which is a source for JPD Bus 10142. It then uses three special microcommands, executed in successive microinstructions, to load S ,Protection Cache Entry 42215. The first microcommand* t *cr loads Protection Cache Entry 42215's TS 24010 with AON 41304 of Logical Descriptor 27116 latched into Descriptor S" Trap 20256; the second loads the object's size into 15 Protection Cache 10234's EXTENT field, and the third E loads the contents of APAME 42106 in the same fashion.

Another microcommand invalidates all Entries 42215 in Protection Cache 10234. This operation, called flushing, .is performed when an object is deactivated or when the 20 current subject changes. The current subject changes whenever a Virtual Processor 612 is unbound from JP 10114, and whenever a Process 610 performs a call to or a return from a Procedure 602 executing in a domain different from that in which the calling Procedure 602 or S' 25 the Procedure 602 being returned to executes in. In the t C TC t cases of the Call and the unbinding of Virtual Processor 612, the cache flush is performed by KOS Call and dispatching microcode; in the case of object -619deactivation, it is performed by a KOS procedure using a special KOS SIN which invokes Cache Flush Microcode.

D. -Processes 1. Synchronization--of--Processes----and Virtual-Ptocessors--6-2 Since Processes 610 and the Virtual Processors 612 to -which they are bound may execute concurrently on CS 10110, KOS must provide means for synchronizing Processes 610 which depend on each other. For example, if Process 610 A cannot proceed until Process 610 B has performed some operation, there must be a mechahism for suspending A's execution until B is finished. Generally speaking, four kinds of synchronization are necessary: One Process 610 must be able to halt and wait S 15s for another Process 610 to finish a task before it proceeds.

f One Process 610 must be able to send another Process 610 a message and wait for a reply S before it proceeds.

S 20 When Processes 610 share a data base, one Process 610 must be able to exclude other S* Processes 610 from the data base until the first Process 610 is finised using the data base.

One Process 610 must be able to interrupt 25 another Process 610, asynchronously cause m tthe second Process 610 to perform some action.

KOS has internal mechanisms for each kind of synchronization, and in addition supplies synchronization mechanisms to EOS. KOS uses the internal mechanisms to 4 p~lll0 -620synchronize Virtual Processors 612 and KOS Processes 610, while EOS uses the mechanisms supplied by KOS to synchronize all other Processes 610. The internal mechanisms are the following: Event counters, Await Entries, and Await Tables. As will be explained in detail below, Event Counters and Await Entries allow one Process 610 to halt and wait for another Process 610 to complete an operation. Event counters and Await Entries are also used to implement process interrupts. Await Entries are organized into Await Tables.

S- Message Queues. Message Queues allow one Process 610 to send a message to another and wait for a reply. Message Queues are implemented with Event Counters and queue data structures.

Locks. Locks allow one Process 610 to exclude other Processes 610 from a data base or a S, 20 segment of code. Locks are implemented with Event Counters and devices called Sequencers.

KOS makes Event Counters, Await Entries, and Message Queues available to EOS. It does hot provide Locks, but it does provide Sequencers, so that EOS can construct its own Locks. The following discussion will define and a jexplain the logical properties of Event Counters, Await Entries, Message Queues, Sequencers, and Locks. Their implementation in the present embodiment will be L_ i -621described along with the implementation of Processes 610 and Virtual Processors 612.

a. Event--Counters--44O-i-T--Await entries-494 and-Avawa-H-aMbfrs -f.-ig--44-8-.-449- Event Counters, Await Entries, and Await Tables are the fundamental components of the KOS Synchronization System. Figure 448 illustrates Event Counters and Await Entries in the present embodiment. Figure 449 gives a simplified representation of Process Event Table 44705, the present embodiment's Await Tables. Turning first to t Figure 448, Event Counter 44801 is an area of memory which contains a value that may only be increased. In one of the present embodiment, Event Counters 44801 for KOS systems which may not page fault are always present in MEM 10112; other Event Counters 44801 are stored in Secondary Storage 10124 unless a Process 610 has referenced them and thereby caused the VMM System to load them into MEM 10112. The value contained in an Event 20 Counter 44801 is termed an Event Counter Value 44802. In the present embodiment, Event Counter 44801 contains 64 bits of data, of which 60 make up Event Counter Value 44802. Event Counter 44801 may be referred to either as a variable or by means of a 128-bit UtD pointer which It..

S 25 contains Event Counter 44801's location. The UID pointer *rt tis termed an Event Counter Name 44803.

Await Entry 44804 is a component of entries in Awash Tables. In the present embodiment, there are two Await i -622- Tables: Process Event Table 44705 and Virtual Processor Await Table (VPAT) 45401. VPAT 45401 is always present in MEM 10112. As already mentioned, Figure 449 illustrates PET 44705. Both PET 44705 and UPAT 45401 will be described in detail later. Each Await Entry 44804 contains an Event Counter Name 44803, an Event Counter Value 44802, and a Back Link 44805 which identifies a Process 610 or a Virtual Processor 612.

Await Entry 44804 thus establishes a relationship between an Event Counter 44801, an Event Counter Value 44802, and a Process 610 or Virtual Processor 612.

Turning now to Figure 449, in the present embodiment, all Await Entries 44804 for user Processes 610 are contained in PET 44705. PET 44705 also contains other information. Figure 449 presents only those parts of PET 44705 which illustrate Await Entries 44804. PET 44705 is structured to allow rapid location of Await Entries 44804 belonging to a specific Event Counter 44801. PET entries (PETEs) 44909 contain links which allow them to be combined into lists in PETE 44705. There are four kinds of lists in PET 44705: Event counter lists: these lists link all PETEs 44909 for Event Counters 44801 whose Event Counter Names 44803 hash to a single value.

25 Await lists: These lists link all PETEs 44909 for Event Counters 44801 which a given Process 610 is awaiting.

4. 4

C,~

1-11,

'I

.4~ 4i P Z -623- Bd r It

I

(I B 4 4* t r *1B4 ii i Interrupt lists: These lists link all PETEs 44909 for Event Counters 44801 which will cause an interrupt to occur for a given Process 610.

The Free list: PETEs 44909 which are not being used in one of the above lists are on a free list.

Each PETE 44909 which is on an await list or an interrupt List is also on an event counter list.

Turning first to the' event counter lists, all PETEs 10 44909 on a given event counter list contain Event Counter Names 44803 which hash to a single value. The value is produced by Hash Function 44901, and then used as an index in PET Hash Table (PETHT) 44903. That entry in PETHT 44903 contains the index in PET 44705 of that PETE 15 44909 which is the head of the event counter list. PETE List 44904 represents one such event counter list. Thus, given an Event Counter Name 44803, KOS can quickly find all Await Entries 44804 belonging to Event Counter 44801.

In the present embodiment, the implementation of 20 Event Counters 44801 and tables with Await Entries 44804 involves both Processes 610 and Virtual Processors 612 to which Processes 610 are bound. As will be explained later, a large number of Event Counters 44801 and Await Entries 44804 belonging to Processes 610 are multiplexed onto a small number of Event Counters 44801 and Await Entries 44804 belonging to the Processes' Virtual Processors 612. Await entries 44804 for Event Counters 44801 belonging to Virtual Processors 612 are contained in VPAT 45401.

L -624b. SynchTonizat-ion--w-ih-Event-eotnt-ers 44.80--and--Await--Entrie s--44-840* The simplest form of Process 610 synchronization provided by KOS uses only Event Counters 44801 and Await Entries 44804. Coordination takes place like this: A Process 610 A requests KOS to perform an Await Operation, to establish one or more Await Entries 44804 and to suspend Process 610 A until one of the Await Entries is satisfied. In requesting the Await Operation-, Process 610 A defines what Event Counters 44801 it is awaiting ,s and what Event Counter Values 44802 these Event Counters 44801 must have for their Await Entries 44804 to be satisfied. After KOS establishes Await Entries 44804, it suspends Process 610 A. While Process 610 A is suspended, other Processes 610 request KOS to perform Advance Operations on the Event Counters 44801 specified in Process 610 A's Await Entries 44804. Each time a «Process 610 requests an Advance Operation on an Event SCounter 44801, KOS increments Event Counter 44801 and checks Event Counter 44801's Await Entries 44804.

Eventually, one Event Counter 44801 satisfies one of -Process 610 A's Await Entries 44804, reaches a value equal to or greater than the Event Counter Value 44802 specified in its Await Entry 44804 for Process 610 tttd 25 A. At this point, KOS allows Process 610 A to resume execution. As Process 610 A resumes execution, it deletes all of its Await Entries 44804.

-625- E. Vir.tual Processos--61-2--5g--45 As previously stated, a Virtual Processor 612 may be logically defined as the means by which a Process 610 gains access to JP 10114. In physical terms, a Virtual Processor is an area of MEM 10112 which contains the information that the KOS microcode which binds Virtual Processors 612 to JP 10114 and unbinds them from JP 10114 requires to perform the binding and unbinding operations. Figure 453 shows a Virtual Processor 612.

The area of MEM 10112 belonging to a Virtual Processor 612 is Virtual Processor 612's Virtual Processor State Block (VPSB) 614. Each Virtual Processor 612 in a CS r 10110 has a VPSB 614. Together, the VPSBs 614 make up VPSB Array 45301'. Within the Virtual Processor management system, each Virtual Processor 612 is known by its VP Number 45304, which is the index of the Virtual Processor 612's VPSB 614 in VPSB Array 45301. Virtual SProcessors 612 are managed by means of lists contained in SMicro VP Lists (MVPL) 45309. Each Virtual Processor 612 has an Entry (MVPLE) 45321 in MVPL 45309, and as Virtual Processor 612 changes state, virtual processor management microcode moves it from one list to another in MVPL 45309.

1 VPSB 614 contains two kinds of information: S 25 information from Process Object 901 belonging to Process 610 which is bound to VPSB 614's Virtual Processor 612, i and information used by the Virtual Processor Management System to manage Virtual Processor 612. The most -626important information from Process Object 901 is the following: Process 610's principal and process UIDs 40401.

AONs 41304 for Proces 610's Stack Objects 44703. (VPSB 614 uses.AONs 41304 because KOS guarantees that AONs 41304 belonging to Stack Objects 44703 will not change as long as a Process 610 is bound to a Virtual Processor 612.) Given AON 41304 of Process 610's SS object 10336, the f'll Virtual Processor Management System can locate that portion of Process 610's state which is moved into registers belonging to JP 10114 when Process 610's Virtual Processor 612 is bound to JP 10114. Similarly, when Virtual Processor 612 is unbound from JP 10114, the virtual processor management system can move the contents of JP 10114 registers into the proper location in SS Object 10336.

a. V-ir tta-1-Pro essa qee n i EOS can perform six operations on Virtual Processors 612: Request VP allows EOS to request a Virtual Processor 612 from KOS.

Release VP allows EOS to return a Virtual 25 Processor 612 to KOS.

S Bind binds a Process 610 to a Virtual Processor 612.

Unbind unbinds a Process 610 from a Virtual Processor 612.

a: '-627- Run allows KOS to bind Process 610's Virtual Processor 612 to JP 10114.

Stop prevents KOS from binding Process 610's Virtual Processor 612 to JP 10114.

As can be seen from the above list of operations, EOS has no direct influence over the actual binding of a Virtual Processor 612 to JP 10114. This operation is performed by a component of KOS microcode called the Dispatcher.

Dispatcher microcode is executed whenever one of four things happens: g Process 610 whose Virtual Processor 612 is currently bound to JP 10114 executes an Await fOperation.

Process 610 whose Virtual Processor 612 is currently bound to JP 10114 executes an Advance Operation which satisfies an Await Entry 44801 for some other Process 610.

Either Interval Timer 25410 or Egg Timer 25412 SL, overflows, causing an Event Signal which invokes Dispatcher microcode.

IOJP Bus 10132 is activated, causing an Event Signal which invokes Dispatcher microcode. IOS 10116 activates IOJP bus 10132 when it loads lr. data into MEM 10112 for JP 10114.

When Dispatcher microcode is invoked by one of these t events, it examines lists in MVPL 45309 to determine which Virtual Processor 612 is to run next. For the purposes of the present discussion, only two lists are important: the running list and the eligible list. In the present embodiment, the running list, headed by I I

I

-628- Running List Head 45321, contains only a single MVPLE 45321, that representing Virtual Processor 612 currently bound to JP 10114. In embodiments with multiple JPs 10114, the running list may have more than one MVPLE 45321. The eligible list, headed by Eligible List Head 45313, contains MVPLEs 45321 representing those Virtual Processors 612 which may be bound to JP 10114. MVPLEs 45321 on the eligible list are ordered by priorities assigned Processes 610 by EOS. Whenever KOS Dispatcher microcode is invoked, it compares the priority of Process r' 610 whose Virtual Processor 612's MVPLE 45321 is on the t t l running list with the priority of Process '610 whose Virtual Processor 612's MVPLE 45321 is at the head of the eligible list. If the latter Process 610 has a higher priority, KOS Dispatcher microcode places MVPLE 45321 belonging to the former Process 610's Virtual Processor 612 on the eligible list and MVPLE 45321 belonging to the latter Process 610's Virtual Processor 612 onto the running list. Dispatcher microcode then swaps Processes 610 by moving state in JP 10114 belonging to the former l Process 610 onto the former Process 610's SS object 10336 and moving JP 10114 state belonging to the latter Process 610 from the latter Process 610's SS object 10336 into JP 10114.

b. Vlrtu1Procesg---rr f l t .454) When a synchronization operation is performed on a l' Process 610, one of the consequences of the operation is a synchronization operation on a Virtual Processor 612.

For example, an Advance Operation which satisfies an E I -629t tr C C Sf( ft It C 4' f CSi 4i *4i 444456 C Ci Await Entry 44804 for a Process 610 causes an Advance Operation which satisfies a second Await Entry 44804 for Process 610's Virtual Processor 612. Similarly, a synchronization operation performed on a Virtual Processor 612 may have a synchronization operation on Virtual Processor 612's Process 610 as a consequence.

For example, if a Virtual Processor 612 performs an operation involving file I/O, Virtual Processor 612's Process 610 must await the completion of the I/O operation.

Figure 454 illustrates the means by which processlevel synchronization operations result in virtual processor-level synchronization operations and viceversa. The discussion first describes the components which transmit process-level synchronization operations to Virtual Processors 612 and the manner in which these components operate. Then it describes the components which transmit virtual processor-level synchronization operations to Processes 610 and the operation of these components.

The first set of components is made up of VPSBA 45301 and VPAT 45401. VPSBA 45301 is shown here with two VPSBs 614: one belonging to a Virtual Processor 612 bound to a user Process 610 and one belonging to a Virtual Processor 612 bound to the KOS Process Manager Process 610. VPAT 45401 is a virtual processor-level table of Await Entries 44804. Each Await Entry 44804 is contained in a VPAT Entry (VPATE) 45403. Each Virtual Processor 612 bound to a Process 610 has a VPAT Chunk 45402 of four VPATEs 45403 in VPAT 45401, and can thus await up to four Event 1 l1 M j r$i -630- Counters 44801 at any given time. The location of a Virtual processor 612's VPAT Chunk 45402 is kept in Virtual Processor 612's VPSB 614. When an Advance Operation satisfies any of the Await Entries 44804 belonging to a Virtual Processor 612, all in Virtual Processor 612's VAT Chunk 45402's Await Entries 44804 are deleted. As in PET 44705, VPATES 45403 containing Await Entries 44804 which are awaiting a given Event Counter 44801 are linked together in a list.

VPATEs 45403 for Virtual Processors 612 bound to user S Processes 610 may contain Await Entries 44804 for user Process 610's Private Event Counter 45405. Private Event Counter 45405 is contained in Process 610's Process Object 901. It is advanced each time an Await Entry 44804 in a PETE 44909 on a PET List belonging to Process 610 is satisfied.

The components operate as follows: When KOS performs an Await Operation on Process 610, it makes Await Entries 44804 in both PET 44705 and VPAT 45401 and puts Process 610's VP 612 on the suspended list in MVPL 45309. As previously described, an Await Entry 44804 in PET 44705 awaits an Event Counter 44801 specified in the Await Operation which created Await Entry 44804. Await Entry Q" t 44804 in VPAT 45401 awaits Process 610's Private Event N 25 Counter 45405. Each time an Await Entry 44804 belonging to Process 610 in PET 44705 is satisfied, Process 610's Private Event Counter 45405 is advanced. The advance of Private Event Counter 45405 satisfies Await Entry 44801 -631for Process 610's Virtual Processor 612 in VPAT 45401, and consequently, KOS deletes Virtual Processor 612's VPATEs 45403 and moves Virtual Processor 612's MVPLE 45321 in MVPL 45309 from the suspended list to the eligible list.

The components which allow a Virtual Processor 612 to transmit a synchronization operation to a Process 610 are the following: Outward Signals Object (OSO) 45409, Multiplexed Outward Signals Event Counter 45407, and PET 44705. OSO 45409 contains Event Counters 44801 which KOS t FU 10120 microcode advances when it perfoms operations which user Processes 610 are awaiting. Event Counters 44801 in OSO 45409 are awaited by Await Entries 44804 in PET 44705. Each time KOS FU 10120 microcode advances an Event Counter 44801 in OSO 45409, it also advances Multiplexed Outward Signals Event Counter 45407. It is awaited by an Await Entry 44804 in VPAT 45401 belonging St to Virtual Processor 612 bound to KOS Process Manager Process 610. When Virtual Processor 62 bound to KOS Process Manager Process 610 is again bound to JP 10114, KOS Process Manager Proces 610 examines all PETEs 44909 belongingtyo the Event Counters 44801 in OSO 45423. If an advance of an Event Counter 44801 in OSO 44801 i satisfied a PETE 44909 Process 610, that Process 610's tcCE 25 Private Event Counter 45405 is advanced as previously described, and Process 610 may again execute.

A user I/O operation illustrates how the components work together. Each user I/O channel has an Event i ~i 2 -632t r St Ie til i ii q* I 1 it II t. ctc ir Counter 44801 in OSO 45409. When a Process 610 performs a user I/O operation on a channel, the EOS I/O routine establish an Await Entry 44804 in the PET 44705 list belonging to Process 610 for the channel's Event Counter 44801 in OSO 45409. When the I/O operation is complete, IOS 10116 places a message to JP 10114 in an area of MEM 10112 and activates IOJP Bus 10132. The activation of IOJP Bus 10132 causes an Event Signal which invokes KOS microcode. The microcode examines the message from IOS 10116 to determine which channel is involved, and then advances Event Counter 44801 for that charnnei in OSO 45409 and Multiplexed Outward Signals Event Counter 45407. The latter advance satisfies an Await Entry 44804 for Process Manager Process 610's Virtual Processor 612 15 in VPAT 45401, and Process Manager Process 610 begins executing. Process Manager Process 610 examines OSO 45409 to determine which Event CQunters 44801 in OSO 45409 have been advanced since the last time process manager Process 610 executed, and when it finds such an Event Counter 44801, it examines the Event Counter Chain in PET 44705 for that Event Counter 44801. If it finds that the advance satisfied any Await Entries 44804 in the Event Counter Chain, it advances Private Event Counter 45405 belonging to Process 610 specified in Await Entry 44804, thereby causing that Process 610 to resume execution as previously described.

F. P.r.acess61 -O Stack-Manipulation This section of the specification for CS 10110 describes the manner in which Process 610's MAS 502 and SS 504 are manipulated. As previously mentioned, in CS s-i -633- 10110, a Process 610's MAS 502 and SS 504 are contained in several objects. In the present embodiment, there are five objects, one for each domain's portion of the Macro Stack (MAS) (MAS Objects 10328 through 10324) and one for the Secure Stack (SS) (SS Object 10336). In other embodiments, a Process 610's MAS 502 may contain objects for user-defined domains as well. Though a Process 610's MAS 502 and SS 504 are contained in many objects, they function as a single logical stack. The division into io several objects is a consequence of two things: the domain component of the protection system, which requires t, that an object referenced by a Procedure 602 have Procedure 602's domain of execution, and the need for a t t t location inaccessible to user programs for micromachine state and state which may be manipulated only by KOS.

Stack manipulati-on takes place under the following circumstances, Sn When a Procedure 602 is invoked or a Return SIN is executed. Procedure 602 invocations are 20 performed by means of a Call SIN. Call causes a transfer of control to the first SIN in the invoked Procedure 602 and the Return SIN causes a transfer of control back to the SIN in the ,invoking Procedure 602 which follows the Call

SIN.

When a non-local Go To SIN is executed. The non-local Go To causes a transfer of control to an arbitrary position in some Procedure 602 which

L

-634was previously invoked by Process 610 and whose invocation has not yet ended.

When a condition arises, an execution of a statement in a program puts the executive Process 610 into a state which requires the execution of a previously established Handler Procedure 602 When a Process 610 is interrupted, when an Interrupt Entry 45718 for Process 610 is satisfied.

c*i Most of the mechanisms involved in stack manipulation are 17 used in Call and Return; these operations are therefore dealt with in detail and the other operations only as they differ from Call and Return. The discussion first introduces Call and Return, then explains the stacks in detail, and finally analyzes Call and Return and the other operations in detail.

Intfroduction-to-Ga--I'--and-Retu-rn As a Process 610 executes a program, it executes Call 4 t 6 and Return SINs. A Call SIN begins an invocation of a Procedure 602, and a Return SIN ends the invocation.

Generally speaking, a Call SIN does the following: It saves the state of Process 610's execution of Procedure 602 which contains the Call SIN.

S 25 Included in this state is the information required to continue Procedure 602's execution after the Call SIN is finished. This portion of the state is termed calling Procedure 602's Macrostate.

2 -635- It creates the state which Process 610 requires to begin execution called Procedure 602.

It transfers control to the first SIN in the called Procedure 602's code.

The Return SIN does the oposite: it releases the state of called Procedure 602 restores the saved state of calling Procedure 602 and transfers control to the SIN in the calling Procedure 602 following the Call SIN. An invocation of a Procedure 602 lasts from the execution of the Call SIN which transfers control to the Procedure 602 j to the execution of the Return SIN which transfers I r t control bac'k to Procedure 602 which contained the Call SIN. The state belonging to a given invocation of a Procedure 602 by a Process 610 is called Procedure 602's invocation state.

While Calls and Returns may be implemented in many different fashions, it is advantageous to implement them using stacks. When a Call creates invocation state for a Procedure 602, that invocation state is added to the top of Process 610's stack. The area of a stack which contains the invocation state of a Procedure 602 is called a frame. Since a called Procedure 602 may call another Procedure 602 and that another, a stack may have any number of frames, each frame containing the j t, a 25 invocation state resulting from the invoation of a Procedure 602 by Process 610, and each frame lasting as long as the invocation it represents. When called i Procedure 602 returns to its caller, the frame upon which

S//

-636it executes is released and the caller resumes execution on its frame. Procedure 602 being currently executed by a Process 610 thus always runs on the top frame of Process 610's MAS 502.

Calls and Returns in CS 10110 behave logically like those in other computer systems using stacks to preserve process 610 state, When a Process 610 executes a Call SIN, the SIN saves as Macrostate the current values of the ABPs, the location of the SIN at which the execution of calling Procedure 602 is to continue, and information such as a pointer to calling Procedure 602's Name Table 10350 and UID 40401 belonging to the S-interpreter object which contains the S-interpreter for Procedure 602's Slanguage. The Call SIN then creates a stack frame for called Procedure 602, obtains the proper ABP values, the location of called Procedure 602's Name Table 10350 and UID 40401 belonging to its S-interpreter object, and begins executing newly-invoked Procedure 602 on the ,rV newly-created stack frame. The Return SIN deletes the stack frame, obtains the ABP values and name interpreter information from the Macrostate saved during the Call SIN. and then transfers control to the SIN at which execution of calling Procedure 602 is to continue.

However- the manner in which Call and Return are implemented is deeply affected by CS 10110's Access Control System- Broadly speaking, there are two classes of Calls and Returns in CS 0110: those which are mediated by KOS and those which are not. In the following discussion, the former class of Calls and Returns are termed Mediated Calls and Returns, and the latter are -637called Neighborhood Calls and Returns. Most Calls and Returns executed by CS 10110 are Neighborhood Calls and Returns; Mediated Calls and Returns are typically executed when a user Procedure 602 calls EOS Procedures 602 and these in turn call KOS Procedures 602. The Mediated Call makes CS 10110 facilities available to user Processes 610 while protecting these CS 10110 facilities from misuse. and therefore generally serves the same purpose as system calls in the present art. As will be seen in the ensuing discussion, Mediated Call requires more CS 10110 overhead than Neighborhood Call. but the extra overhead is less than that generally required by system calls in the present art.

Mediated Calls and Returns involve S-interpreter.

Namespace, and KOS microcode. S-interpreter and Namespace microcode interprot the Names involved in the call and only modifies those portions of Macrostate *accessible to the S-interpreter. The remaining Macrostate is modified by KOS microroutines invoked in the course of the Call SIN. A Mediated Call may be made to any Procedure 602 contained in an object to which Process 610's subject has Execute Access at the time the invocation occurs. Mediated Calls and Returns must be made in the following situations: When called Procedure 602 has a different Procedure Environment Descriptor (PED) 30303 from that used by calling Procedure 602. Such Calls are termed Cross-PED Calls.

4 r L~* -f38r 'I C

U.

I I U. I When called Procedure 602 is in a different Procedure Object 608 from calling Procedure 602. Such Calls are termed Cross-Procedure Object Calls.

When called Procedure 602's Procedure Object 608 has a different Domain of Execution (DOE) Attribute from that of calling Procedure 602's Procedure Object 608. and therefore must place its Invocation State on a different MAS object from that used by calling Procedure .602. Such Calls are termed Cross-Domain Calls.

In all of the above Calls. the information required to complete the Call is not available to the S-interpreter.

and consequently. KOS mediation is required to complete 15 the Call- Neighborhood Calls and Returns only modify two components of Macrostate: the pointer to the current SIN and the FP ABP. Both of these components are available to the S-interpreter as long as called Procedure 602 hP, the same PED 30303. uses the same Name Tabe 10350 and S-interpreter or the calling Procedure 602 and has Names with the same syllable size as calling Procedure 602. The Call and Return SINs are specific to each S-language. but they resemble each other in their general behavior. The following discussion will deal exclusively with this general behavior, and will concentrate on Mediated Calls and Returns. The discussion first describes MAS 502 and SS 504 belonging to a Process 610 and those parts of Procedure Object 608 involved in Calls

L

-639and Returns, and then describes the implementation of Calls and Returns.

2- Macro Stacks MASiL502 'Fig- 46-7-) Figure 467 gives an overview of an object belonging to a Process 610's MAS 502. The description of this Figure will be followed by descriptions of other Figures containing detailed representations of portions of MAS objects.

At a minimum. MAS Object 46703 comprises KOS MAS

S'"

1 10 Header 10410 together with Unused Storage 46727 reserved for the other elements comprising MAS Object 46703. If Process 610 has not yet returned from an invocation of a Procedure 602 contained in a Procedure Object 608 whose DOE is that required for access to MAS Object 46703. MAS object 46703 further comprises a Stack Base 46703 and at least one MAS Frame 46709.

t Each MAS Frame 46709 represents one mediated C invocation of a Procedure 602 contained in a Procedure Objece 608 with the DOE attribute required by MAS 46703.

20 and may in addition represent neighborhood invocations of Procedures 602 which share that Procedure 602's Procedure Object 608. The topmost MAS Frame 46709 represents the most recent group of invocations of Procedures 602 with the DOE attribute required by MAS Object 46703. and the bottom MAS Frame 46709 the earliest group of invocations from which Process 610 has not yet returned. Frames for invocations of Procedures 602 with other domains of execution are contained in other MAS Objects 46703. As will be explained in detail below. MAS Frames 46709 in different MAS objects 46703 are linked by pointers.

e^a~i. ii -640- MAS Domain Stack Base 46703 has two main parts: KOS MAS Header 10410- which contains information used by KOS microcode which manipulates MAS Object 46703. and Perdomain Information 46707, which contains information about 46703's domain and static information, i.e., information which lasts longer than an invocation, used by Procedures 602 with MAS Frames 46709 on MAS Object 46703- MAS Frame 46709 also has two main parts, a KOS Frame Header 10414- which contains information used by 10 KOS to manipulate Frame 46709. and S-interpreter Portion *46713. which contains information available to the Sinterpreter when it executes the group of Procedures 602 whose invocatios are represented by Frame 46709.

When making Calls and Returns, the S-interpreter and J 15 KOS microcode use a group of pointers to locations in MAS Object 46703- These pointers comprise the following: S, MAS Object UID 46715. the UID 40401 of AS Object 46703.

First Frame Offset (FFO) 46719. which locates S: 20 the beginning of KOS Frame Header 10414 belonging to the first MAS Frame 46709 in MAS Object 46703.

Frame Header Pointer (FHP) 46702. which locates Ithe beginning of the topmost KOS Frame Header S 25 10414 in MAS Object 46703.

Stack Top Offset (STO) 46704. a 32-bit offset from Stack UID 46715 which marks the first bit in Unused Storage 46727.

I

L r -641- As will be seen presently. all of these pointers are contained in fields in KOS MAS Header 46705.

a.a. MAS Base 10_41_(Fig. 468) Figure 468 is a detailed representation of MAS Domain Stack Base 10410- Turning first to the detailed representation of KOS MAS Header 46705 contained therein, there are the following fields: Format Information Field 46801. containing information about the format of KOS MAS Header 46705.

Flags Field 46803. Of these flags, only one is of interest to the present discussion: Domain Active Flag 46804- This flag is set to TRUE when Process 610 to which MAS Object 46703 belongs is executing the invocation of Procedure 602 whose invocation record makes up the topmost MAS Frame 46709 contained in MAS Object 46703 to C| V which KOS MAS Header 46705 belongs.

St PFO Field 46805: All MAS Headers 46705 and Frame Headers 46709 have fields containing offsets Slocating the previous and following headers in MAS Object 46703- In a Stack Header 46705.

there is no previous header, and this field is S" set to 0.

FFO Field 46805: The field locating the following header. In a Stack Header 46705. this field contains FFO 46719. since the next header is the first Frame Header in MAS Object 46703.

Lf -642- STO Field 46807: the field containing STO offset 46704.

Process ID Field 46809: UTD 40401 belonging to Process Object 901 for Process 610 to which MAS Object 46703 belongs.

Domain Environment Information Pointer Field 46811: The pointer contained in the field locates an area which contains domain-specific information. In the present embodiment. the 10 area is part of MAS Stack Base 10410; however, in other embodiments, it may be contained in a t separate object.

Signaller Pointer Field 46813: The pointer contained in the field locates a Procedure 602 which KOS invokes when a Process 610's execution causes a condition to arise while it is C V tf executing in the domain to which MAS object St, 46703 belongs.

AAT Pointer Field 30211: The pointer in Field 30211 locates AAT 30201 for MAS Object 46703.

AAT 30201 is described in detail in Chapter 3.

Frame Label Sequencer Field 46819: This field i9*9 contains a Sequencer 45102. Sequencer 45102 is used to generate labels used to locate MAS Frames 46709 when a non-local GOTO is executed.

Turning now to the detailed representation of Domain Environment Information 46821 located by Domain

Y

-643- Environment Information Pointer Field 46811. there are the following fields: KOS Format Information Field 46823.

Flags Field 46825. containing the following flags: Pending Interrupt Flag 46827, set to TRUE when Process 610 has an interrupt pending for the domain to which MAS Object 46703 belongs.

Domain Dead Flag 46829, set to TRUE when Process 610 can no longer execute SProcedures 602 with domains of execution equal to that to which MAS Object 46703 belongs.

Invoke Verify on Entry Flag 46833 and Invoke Verify on Exit Flag 46835. The Sformer flag is set to TRUE when KOS is to Sinvoke a Procedure 602 which checks the domain's data bases before a Procedure 602 20 is allowed to execute on the domain's MAS Object 46703; the latter is set to TRUE when KOS is to invoke such a Procedure 602 i on exit from a Procedure 602 with the Sfor tedomain as its DOE.

Default Handler Non-null Flag 46835 is set to TRUE when there is a default clean-up handler for the domain. Clean-up handlers are described later. F8 al fo te d n -644- Interrupt Mask Field 46839 determines what interrupts set for Process 610 in MAS object 46703's domain will be honored.

Domain UID Field 46841 contains UID 40401 for the domain to which MAS Object 46703 belongs.

Fields 46843 through 46849 are pointers to Procedures 602 or tables of pointers to Procedures 602. The Procedures 602 so located handle situations which arise as MASs 502 are 1, 0 manipulated. The use of these fields will Sbecome clear as the operations which require their use are explained.

t t b.b. ERr-domainData Area46_85 (Fig. 468) Per-domain Data Area 46853 contains data which cannot be kept in MAS Frames 46709 belonging to invocations of Procedures 602 executing in MAS Object 46703's domain, but which must be available to these invocations. Per- Domain Data Area 46853 has two components: Storage Area 46854 and AAT 30201. Storage Area 46854 contains static S* data used by Procedures 602 with invocations on MAS Object46703 and data used by S-interpreters which are used by such Procedures 602. Associated Address Table (AAT) 30201 is used to locate data in Storage Area 46854. A detailed discussion of AAT 30201 is contained in Chapter 3.

Two kinds of data is stored in Storage Area 46854: i static data and S-interpreter data.

-645- Static data is stored in Static Data Block 46863.

Static Data Block 46863 comprises two parts: Linkage Pointers 46865 and Static Data Storage 46867. Linkage Pointers 46865 are pointers to static data not contained in Static Data Storage 46867. for example, data which lasts longer than Process 610. and pointers to External Procedures 602 which the Procedure 602 to which Static Data Storage 46867 belongs invokes. Static Data Storage 46867 contains storage for static data used by the 10 Procedure 602 which does not last longer than Process 610 executing the Procedure 602r t S-interpreter data is data required by S-interpreters used by Procedures 602 executing on MAS object 46703.

The S-interpreter data is stored in S-interpreter Environment Block (SEB) 46864. which, like Static Data Block 46864. is located via AAT 30201- The contents of S, SEB 46864 depend on the S-interpreter.

c.c. MAS Frame- 46709 Detail ;Eig, 20 Figure 469 represents a typical frame in MAS Object 46703. Each MAS Frame 46709 contains a Mediated Frame 46947 produced by a Mediated Call of a Procedure 602 contained in a Procedure Object 608 whose DOE attribute tJB ru is the one required for execution on MAS object 46703.

Mediated Frame 46947 may be followed by Neighborhood Frames 46945 produced by Neighborhood Calls of Procedures 602- Mediated Frame 46947 has two parts, a KOS Frame Header 10414 which is manipulated by KOS microcode, and an S-interpreter Portion which is manipulated by Sinterpreter and Namespace microcode. Neighborhood Frames -646- 46945 have no KOS Frame Headers 10414. As will become clear upon closer examination of Figure 469. Mediated Frames 46947 in the present embodiment contain no Macrostate. In the present embodiment. Macrostate for these frames is kept on SS Object 10336; however in other embodiments, Macrostate may be stored in Mediated Frames 46947. Neighborhood Frames 46945 contain those portions of the macrostate which may be manipulated by Neighborhood Call; the location of this macrostate S 10 depends on the Neighborhood Call SIN- Turning now to KOS Frame Header 10414. there are the following fields: KOS Format Information Field 46901. containing information about MAS Frame 46709's format.

Flags Field 46902. This field contains the following flags: Result of Cross-domain Call Flag 46903.

This Flag is TRUE if MAS Frame 46709 which B precedes this MAS Frame 46709 is in another MAS Object 46703.

Is Signaller Flag 46905. This flag is TRUE if this MAS Frame 46709 was created by the invocation of a Signaller Procedure 602.

Do Not Return Flag 46907: This flag is TRUE S* 25 if Process 610 is not to return to the invocation for which this MAS Frame 46709 was created.

R'

-647- *9g* 4 9494

I

*r t4 #9 944- 9 9 999499t 9 9 9 9 9 #999 4I i 4.

I

4 .4 Flags 46909 through 46915 indicate whether various lists used in condition handling and non-local GOTOs are present in the MAS Frame 46709.

Previous Frame Offset Field 46917, Next Frame Offset Field 46919. and Frame Top Offset Field 46921 are offsets which give the location where Header 10414 for the previous MAS Frame 46709 in MAS Object 46703 begins, the location where the header for the next MAS Frame 46709 in MAS Object 46703 begins, and the location of the first bit beyond the top of MAS Frame 46709 respectively.

Fields 46923 through 46927 are offsets which locate lists in S-interpreter portion 46713 of Frame 46709. KOS establishes such lists to handle conditions and non-local GOTOs. Their use will be explained in detail under those headings.

Fields 46929 and 46933 contain information about Procedure 602 whose invocation is represented by MAS Frame 46709. Field 46929 contains the number of arguments required by Procedure 602.

and Field 46933 contains a resolvable pointer to Procedure 602's PED 30303. Both these fields are used primarily for debugging.

Dynamic Back Pointer Field 46931 contains a resolvable pointer to the preceding MAS Frame 46709 belonging to Process 610's MAS 502 when that MAS Frame 46709 is contained in a different 44 4 44 444.

4*.

4* 4 9 149414r 4 4- A 1 -648- MAS Object 46703. In this case. Flag Field 46903 is set to TRUE. When the preceding MAS Frame 46709 is contained in the same MAS object 46703. field 46931 contains a pointer with a null UID 40401 and Flag Field 46903 is set to

FALSE.

Frame Label Field 46935 is for a Frame Label produced when a non-local GOTO is established which transfers control to the invocation 10 represented by MAS Frame 46709. The label is generated by Frame Label Sequencer 46819 in KOS o MAS Header 10410.

S-interpreter Portion 46713 of MAS Frame 46709 comprises those portions of MAS Frame 46709 which are under control of the S-interpreter- S-interpreter Portion 46713 in turn comprises two main subdivisions: those parts S|belonging to Mediated Frame 46947 and those belonging to 4 Neighborhood Frames 46945.

t* The exact form of S-interpreter portion 46949 of KOS Frame 46947 and of S-interpreter Frames 46945 depends on the Call SIN which created the frame in question.

However, all Neighborhood Frames 46945 and S-interpreter I Portions 46949 of Mediated Frames 46947 have the same arrangements for storing Linkage Pointers 10416 and local S* 25 data in the frame. Linkage Pointers 10416 are pointers to the locations of actual arguments used in the invocation and Local Storage 10420 contains data which exists only during the invocation. In all Mediated Frames 46947 and Neighborhood Frames 46945. Linkage Pointers 10416 precede Local Storage 10420. Furthermore.

-649when a Mediated Frame 46947 or a Neighborhood Frame 46945 is the topmost frame of Process 610's MAS. i.e. when Process 610 is executing on that frame, the FP always points to the beginning of Local Storage 10420. and the beginning of Linkage Pointers 10416 is always at a known displacement from FP. References to Linkage Pointers 10416 may therefore be expressed as negative offsets from FP, and references to Local Storage 10420 as positive offsets.

1 10 In addition, S-interpreter Portion 46713 may contain lists of information used by KOS to execute non-local SGOTOs and conditions, as well as S-interpreter frames for non-mediated calls. The lists of information used by KOS SC are contained in List Area 46943. The exact location of S 15 List Area 46943 is determined by the compiler which generates the SINs and Name Table for the Procedure 602 whose invocation is represented by Mediated Frame 46947.

When Procedure 602's source text contains statements requiring storage in List Area 46943. the compiler generates SINs which place the. required amount of storage in Local Storage 10420. KOS routines then build lists in Area 46943. and place the offsets of the heads of the lists in Fields 46923, 46925. or 46927, depending on the I kind of list. The lists and their uses are described in detail later.

3. SS 504A (ig. _410 Figure 470 presents an overview of SS 504 belonging to a Process 610. SS 504 is contained in SS Object 0 i.

-650- 10336. SS Object 10336 is manipulated only by KOS microcode routines. Neither Procedures 602 being executed by Process 610 nor S-interpreter or Namespace microcode may access information contained in SS Object 10336.

SS Object 10336 comprises two main components, SS Base 47001 and SS Frames 47003. Turning first to the general structure of SS Frames 47003. each time a Process 610 executes a Mediated Call. KOS microcode creates a new 10 SS Frame 47003 on SS Object 10336 belonging to Process 610. and each time a Process 610 executes a Mediated Return. KOS microcode removes the current top SS Frame 47003 from SS Object 10336. There is thus one SS Frame 47003 on SS Object 10336 belonging to a Process 610 for each Mediated Frame 46947 on Process 610's MAS 502.

SS Frames 47003 comprise two kinds of frames: Ordinary Frames 10510 and Cross-domain Frames 47039.

Cross-domain Frames 47039 are created whenever Process 610 executes a Cross-domain Call; for all other Mediated 20 Calls, Ordinary Frames 10510 are created. Cross-domain Frames 47039 divide SS Frames 47003 into Groups 47037 of SS Frames 47003 belonging to sequences of invocations in a single domain. The first SS Frame 47003 in a Group 47037 is a Cross-domain Frame 47039 for the invocation which entered the domain, and the remainder of the SS Frames 47003 are Ordinary Frames 1051,0 for a sequence of invocations in that domain. These groups of SS Frames 47003 correspond to groups of Mediated Frames 46947 in a single'MAS Object 46703.

-651a.a. SBAe_4700QL_CEL.. .421) SS Base 47001 comprises four main parts: SS Header 10512- Process Microstate 47017, Storage Area 47033 for JP 10114 register contents, and Initialization Frame Header 47035. Secure Stack Header 10512 contains the following information: Fields 47001 and 47009 contain flag and format information; the exact contents of these fields are unimportant to the present discussion.

S 10 Previous Frame Offset Value Field 47011 is a o standard field in headers in SS Object 10336; here. it is set to 0. since there is no previous frame.

Secure Stack First Frame Offset Field 47013 contains the offset of the first SS Frame 47039 in SS object 10336. Initialization Frame Header 47035.

Process UID field 47015 contains UTD 40401 of Process 610 to which SS Object 10336 belongs.

20 Number of Cross Domain Frames Field 47016 contains the number of Cross-domain Frames 47039 in SS Object 10336.

Process Microstate 47017 contains information used by KOS microcode when it executes Process 610 to which SS Object 10336 belongs. Fields 47019. 47021. and 47022 contain the offsets of locations in SS Object 10336. Field 47019 contains the value of SSTO. the location of the first free bit in SS Object 10336; Field 47021 contains the -652value of SSFO, the location of the topmost frame in SS object 10336; Field 47022. finally, contains the value of XDFO. the location of the topmost Cross-domain Frame 47039 in SS Object 10336. All of these locations are marked in Figure 470.

Other fields of interest in Process Microstate 47017 comprise the following: Offsets in Storage Area Field 47023 contains offsets of locations in Storage Area 47033 .of SS Object 10336; Domain Number Field 47025 contains t10 the domain number for the DOE of Procedure 602 currently being executed by Process 610- The relationhip between I domain UIDs and domain numbers is explained in the discussion of domains. VPAT Offset Field 47027 contains C the offset in VPAT 45401 of VPAT Chunk 45402 belonging to Virtual Processor 612 to which Process 610 is bound.

Signal Pointer Field 47029 contains a resolved pointer to the Signaller (a Procedure 602 used in condition

S;

6 handling) belonging to the domain specified by Domain Number Field 47025. and Trace Information Field 47031 20 contains a resolved pointer to that domain's Trace Table, described later.

Storage Area for JP 10114 register Contents 47033 is i used when a Virtual Processor 612 must be removed from JP 10114. When this occurs, either because Virtual Processor 612 is unbound from JP 10114. because CS 10110 is being halted, or because CS 10110 has failed, the K contents of JP 10114 registers which contain information specific to Virtual Processor 612 are copied into Storage

F

n C--i:ILC: i- -653- Area 47033- When Virtual Processor 612 is returned to JP 10114. these register contents are loaded back into the JP 10114 registers from whence they came. Initialization Frame Header 47035- finally, is a dummy frame header which is used in the creation of SS Object 10336.

b.b. SS Frames 4700- (Figa. 471 Commencing the discussion of SS Frames 47039 and 10510. Figure 471 illustrates these structures in detail.

Ordinary SS Frame 10510 comprises three main divisions: 10 Ordinary SS Frame Header 10514. Macrostate 10516. and t tt Microstate 10520. Ordinary SS Frame Header 10514 contains information used by KOS microcode to manipulate Ordinary SS Frame 10510 to which Header 10514 belongs.

Macrostate 10516 contains the values of the ABPs for the frame's mediated invocation and other information required to resume execution of the invocation.

Microstate 10520 contains micromachine state from FU 10120 and EU 10122 registers. The amount of micromachine state depends on the circumstances; in the present 20 embodiment, some micromachine state is saved on all Mediated Calls; furthermore, if a Process 610 executes a microcode-to-software Call. the micromachine state that existed at the time of the call is saved; finally.

Microstate 10520 belonging to the topmost SS Frame 47003 25 may contain information which was transferred from FU 10120 GRF registers 10354 or EU 10122 register and stack mechanism 10216 when their capacity was exceeded. For details about this portion of Microstate 10520. see the discussion of the FU 10120 micromachine in Chapter 2.

4 i'" 2 1 I i 1 -654- The discussion of SS Object 10336 continues with details concerning SS Header 10514 and Macrostate 05163.

a.a.a. Ordinary- SS--Frame--Headers 1 514_(ig.- 4711 Fields of interest in Ordinary Secure Stack Frame Header 10514 are the following: Format Information 47103. which identifies the format of Header 10514.

Flags Field 47105. which contains one flag of interest in this discussion: Frame Type Flag 47107: in Ordinary SS Frames 10510. this field is set to FALSE.

Offset Fields 47109 through 47113: Field 47109 contains the offset of the previous SS Frame 47039 or 10510. Field 47111 contains the offset of the following SS Frame 47039 or 10510. and Field 47113 contains the offset of the last SS Frame 47039 or 10510 preceding the next Cross- S- domain Frame 47039.

Field 47117 contains the current domain number for the domain in which the mediated invocation S, r represent by SS Frame 47039 or 10510 is executing.

Field 47119 contains the offset of the preceding Cross-domain Frame 47039.

Field 47121 contains offsets for important locations in Microstate 10520.

|J tn reresnt-b SSFram 4709 o 1050 i pr 4 4. tr .r UI ra -655b.b.b. DetailedStructure of Macrostate 10516 figa 4711 These fields are of interest in Macrostate 10516: Syllable Size Field 47125 contains the value of K, the size of the Names in the SINs belonging to Procedure 602 which the invocation is executing.

End of Name Table Field 47127 contains the location of the last Name in Name Table 10350 belonging to Procedure 602 which the invocation is executing.

Fields 47129 through 47143 are resolved pointers to locations in Procedure Object 901 containing Procedure 602 being executed by the invocation and resolved pointers to locations containing data being used by Procedure 602. Field 47129 contains a pointer to Procedure 602's PED 30303; if Procedure 602 is an External Procedure 602.

Field 47131 contains a pointer to Procedure 602's entry in Gates 10340; Field 47135 contains the UID-offset value of FP for the invocation; Field 47135 contains a pointer to SEB 46864 used by Procedure 602's S-interpreter. Field 47137 contains the UID-offset value of SDP and Field 47139 contains that of PBP. SIP Field 47141 contains a pointer to Procedure 602's Sinterpreter object, and NTP, finally, is a pointer to Procedure 602's Name Table 10350.

11~11120 I 44 4.44.

4. 4.

4a4 ,'4 i"I -656- Field 47145 contains the PC for the SIN which is to be executed on return from the mediated invocation to which SS Frame 47003 belongs.

c.C.c. eross-daomain--s.-Pmes Cross-domain SS Frames 47039 differ from Ordinary SS Frames 10510 in two respects: they have an additional component, Cross-domain State 10513. and fields in Cross domain Frame Header 47157 have different meanings from those in Ordinary Frame Header 10514.

Cross-domain State 10513 contains information which KOS Call microcode uses to verify that a return to a Procedure 602 whose DOE differs from that of Procedure 602 whose invocation has ended is returning to the proper domain. Fields of interest in Cross-domain State 10513 include GOTO Tag 47155. used for non-local GOTOs which Scross domains, Stack Top Pointer Value 47153, which gives Se the location of the first free bit in the new domain's SMAS Object 46703. and Frame Header Pointer Value 47151.

A 20 which contains the location of the topmost Mediated Frame Header 46709 in new MAS Object 46703.

There are three fields in Cross-domain Frame Header 4: 7157 which differ from those in Ordinary SS Frame Header 47101. These fields are Flag Field 47107, which in Cross-domain Frame Header 47157 always has the value TRUE, Preceding Cross-domain Frame Offset Field 47161.

which contains the offset of preceding Cross-domain Frame 47039 in SS Object 10336. and Next Cross-domain Frame Offset Field 47159, which contains the location of the a. I are Flag Field 4710,wichi TRUE Prcedig Coss-oman Frme ffse Fild 4161 whic cotain th offet f pecedng rossdo^in Fame r in SS c 10336-.

e j -657next Cross-domain Frame 47039. These last two fields occupy the same locations as Fields 47111 and 47109 respectively in Ordinary SS Frame Header 10514.

As will be noted from the above description of SS Frames 47003. Secure Stack Object 10336 in the present embodiment contains three kinds of information: macrostate. cross-domain state, and microstate. In other embodiments, the information in SS object 10336 may be stored in separate stack structures, for example, cc.. 10 separate microstate and cross-domain stacks, or information presently stored in MAS Objects 46703 may be stored in SS Object 10336. and vice-versa.

4. Paotions-of-P-rocedure-eb+eet--j-6 The information which Process 610 requires to construct new frames on its MAS Objects 46703 and SS Object 10336 and to transfer control to invoked Procedure 602 is contained in invoked Procedure 602's Procedure Object 608. Figure 472 is an overview of Procedure 20 Object 608 showing the information used in a Call.

Figure 472 expands information contained in Figures 103 and 303; fields that appear in those Figures have the t: names and numbers used there.

Beginning with Procedure Object Header 10336, this area contains two items of information used in Calls: an offset in Field 47201 giving the location of Argument Information Array 10352 in Procedure Object 608 and a value in Field 47203 specifying the number of gates in Procedure Object 608. Gates allow the invocation of 1.

Ur; *14 -658- Nt External Procedures 602. that is, Procedures 602 which may be invoked by Procedures 602 contained in other Procedure Objects 608. Procedure Object 608's gates are contained in External Entry Descriptor Area 10340. There are two kinds of gates: those for Procedures 602 contained in Procedure Object 608. and those for Procedures 602 contained in other Procedure Objects 608, but callable via Procedure Object 608. Gates for Procedures 602 contained in Procedure Object 608 are 0' 0 termed Local Gates 47205. Local Gates 47205 contain t Internal Entry Offset (IEO) Field 47207 which contains ?t* 8 the offset in Procedure Object 608 of Entry Descriptor 47227 for Procedure 602- If Procedure 602 is not contained in Procedure object 472. its gate is a Link Gate 47206. Link Gates 47206 contain Binder Area Pointer (BAP) 'Fields 47208. A BAP Field 47208 contains the locations of an area in Binder Area 30323 which in turn contains a pointer to a Gate in another Procedure Object 608. The pointer in Binder Area 30323 may be either 20 resolved or unresolved. If Procedure 602 is contained in that Procedure Object 608. the Gate is a Local Gate 47205; otherwise, it is another Link Gate 47206.

Procedure Environment Descriptors (PEDS) 10348 "contains PEDs 30303 for Procedures 602 contained in Procedure Object 608. Most of the macrostate information for a Procedure 602 may be found in its PED 30303. PED 30303 has already been described, but for ease of understanding, its contents are reviewed here.

K Field 30305 contains the size of Procedure 602's Names.

-659- Largest Name (LN) Field 30307 contains the i Beginning with Procedure Object Header 10336. this area contains two items of information used in Calls: an Offset in Field 47201 giving the location of Argument Information Array 10352 in Procedure Object 608 and a value in Field 47203 specifying the number of gates in Procedure Object 608. Gates allow the invocation of External Procedures 602. that is, Procedures 602 which t may be invoked by Procedures 602 contained in-other "t 10 Procedure Objects 6nter to Static Data Block 46863.

Thus, for that invocation of Procedure 602. on invocation, the SDP ABP is derived via SDPP field 30313.

PBP Field 30315 is the pointer from which the current PC is calculated. When Procedure 602 is invoked, this value becomes the PBP ABP.

S-interpreter Environment Prototype Pointer (SEPP) Field 30316 contains the location of SEB CIt Prototype Field 30317. When Procedure 602 is invoked, Field 30316 locates SEB 46864 via AAT 20 30201 in the same manner as SDPP field 30313 locates the invocation's static data.

A Procedure 602's FED 30303 may be located from its Internal Entry Descriptor 47227. A FED 30303 may be t shared by several Procedures 602. Of course, in this case, the values contained in shared PED 30303 are the same for all Procedures 602 sharing it. As will be explained in detail later. in the present embodiment, if a calling Procedure 602 does not share a PED 30303 with called Ptocedure 602. the Call must be mediated. A

C-

L 4 .1 2 -660calling Procedure 602 may make a Neighborhood Call only to Procedures 602 with which it shares a PED 30303.

The next portion of Procedure Object 608 which is of interest is Internal Entry Descriptors 10342. Each Procedure 602 contained in Procedure Object 608 has an Entry Descriptor 47227. Entry Descriptor 47227 contains four fields of interest: PBP Offset Field 47229 contains the offset from PBP at which the first SIN in Procedure 602's 10 code is located.

Flags Field 47230 contains flags which are checked when Procedure 602 is invoked. Four flags are of interest: Argument Information Array Present Flag 47235. which is set to TRUE if Procedure 602 has entries in Argument Information Array 10352.

SEB Flag 47237 is set to TRUE if SEPP 47225 is non-null, if Procedure 602 has a t 20 SEB 46864 for its S-interpreter.

Do Not Check Access Flag 47239 is set to TRUE if KOS Call microcode is not to perform protection checking on the actual ,tt arguments used to invoke Procedure 602.

25 PED Offset Field 47231 contains the offset of Procedure 602's PED 30303 from the beginning of Procedure Object 608.

i 'i -661- Frame Size Field 47233 contains the initial size of the Local Storage Portion 10420 of MAS Frame 46709 for an invocation of Procedure 602- Other areas of interest for Calls are SEB Prototype Area 47241. Static Data Area Prototype 30317. Binder Area 30323. and Argument Information Array 10352. SEB Prototype type Area 47241 and Static Data Area Prototype 30315 contain information used to create an SEB 46864 and Static Data Block 46863 respectively for Procedure 602.

10 These areas are created on a per-MAS Object 46703 basis.

The first time that a Process 610 executes a Procedure 602 in a domain, SEB 46864 and Static Data Block 46863 required for Procedure 602 are created either in MAS Object 46703 belonging to the domain or in another object accessible from MAS Object 46703. SEB 46864 and Static Data Block 46863 then remain as long as MAS Object 46703 exists.

Static Data Prototype 30317 contains two kinds of information: Static Data Links 30319 and Static Data j 20 Initialization Information 30321. Static Data Links 30319 contain locations in Binder Area 30323. which in turn contains pointers which may be resolved to yield the locations of data or .External Procedures 602. When a Static Data Block 46863 is created for a Procedure 602.

the information in Binder Area 30323 is used to create Linkage Pointers 46865. Static Data Initialization Information 30321 contains information required to create ,1 and initialize static data in Static Data Storage 46867. I -662- As mentioned in the discussions of Link Gates 47206 and Static Data Links 30319. Binder Area 30323 contains pointers which may be resolved as described in Chapter 3 to yield locations of data and External Procedures 602.

Argument Information Array (AIA) 10352 contains information used by KOS Call microcode to check whether the subject which is invoking Procedure 602 has access to the actual arguments used in the invocation which allows the uses made of the arguments in Procedure 602. This 10 so-called "Trojan horse check" is necessary because a Call may change the domain component of a subject. Thus, Sa subject which is lacking access of a specific kind to a r, data item could gain that access by passing the data item as an argument to a Procedure 602 whose DOE gives it access rights that the calling subject itself lacks.

Each Local Gate 47205 in Procedure Object 608 has an t V V t element in AIA 10352. Each of these Argument Information Array Elements (AIAEs) 60845 has fields indicating the following: 20 The minimum number of arguments required to invoke Procedure 602 to which Local Gate 47205 belongs, in Field 47247.

t The maximum number of arguments which may be r used to invoke Procedure 602. in Field 47249.

The access rights that the invoking subject must have to the actual arguments in order to invoke Procedure 602. in Fied 47251. r hh.

41 -663- Field 47251 is itself an array which specifies the kinds of access that the invoking subject must have to the actual arguments it uses to invoke Procedure 602. Each formal argument for Procedure 602 has an Access Mode Array Entry (AMAE) 47255. The order of the AMAEs 47255 corresponds to the order of Procedure 602's formal arguments. The first formal argument has the first AMAE 47255. the second the second, and so forth. An AMAE 47253 is four bits long. There are two forms of AMAE 47253: 10 Primitive Access Form 47255 and Extended Access Form 47257. In the former form. the leftmost bit is set to S' 0. The three remaining bits specify read. write, and I execute access. If a bit is on, the subject performing the invocation must have that kind of primitive access to the object containing the data item used as an actual for the formal argument corresponding ,to that AMAE 47253. In the Extended Access Form 47257, the leftmost bit is set to 1 and the remaining bits are defined to represent extended access required for Procedure 602. The 20 definition of these bits variesfrom Procedure 602 to Procedure 602.

*1t 1*C 1 Execution- o MeCals Having described the portions of MAS Object 46703. SS Object 10336. and Procedure Object 608 which are involved in Calls, the discussion turns to the description of the Mediated Call Operation. First, there is presented an overview of the Mediated Call S'IN. and then the implementation of Mediated Calls in the present -664embodiment is discussed. beginning with a simple Mediated Call and continuing with Cross-Procedure Object Calls and Cross Domain Calls. The discussion closes with a description of software-to-microcode Calls.

a.a. Mediated-ea- S-INs While the exact form of a Mediated Call SIN is Slanguage specific, all Mediated Call SINs must contain four items of information: The SOP for the operation.

A Name that evaluates to a pointer to the I Procedure 602 to be invoked by the SIN.

A literal (constant) specifying the number of actual arguments used in the invocation.

A list of Names which evaluate to pointers to the actual arguments used in the invocation.

If Procedure 602 requires no arguments, the literal will be 0 and the list of Names representing the actual arguments will be empty.

In the present embodiment, Mediated Call and Return 20 SINs are used whenever called Procedure 602 has a different PED 30303 from calling Procedure 602. In this case, the Call must save and recalculate macrostate other 4 than FP and PC, and mediation by KOS Call microcode is required. The manner in which I~iS Call microcode mediates the Call depends on whether the Call is a simple Mediated Call. a Cross-procedure Object Call. or a Cross- Domain Call- LI. A i o a w el t o r -665b.b. Simple- Mediated- Calls-ig--27.

468. 469 ;2470-.- 4 71-. 472 When the Mediated Call SIN is executed. S-interpreter microcode first evaluates the Name which represents the location of the called Procedure 602. The Name may evaluate to a pointer to a Gate 47205 or 4707 in another Procedure Object 608 or to a pointer to an Entry Descriptor 47227 in the present Procedure Object 608.

When the Name has been evaluated. S-interpreter Call microcode invokes KOS Call microcode, using the evaluated Name as an argument. This microcode first fills in Macrostate Fields 10516. left empty until now. in the current invocation's SS Frame 47003. The microcode obtains the values for these fields from registers in FU 10120 where they are maintained while Virtual Processor 612 of Process 610 which is executing the Mediated Call is bound to JP 10114.

SThe next step to determine whether the pointer which XOS Call microcode received from S-interpeter Call 20 microcode is a pointer to an External Procedure. To make this determination, KOS Call microcode compares the pointer's AON 41304 with that of Procedure Object 608 for Procedure 602 making the Call. If they are different, the Call is a Cross-Procedure Object Call. described below. In the case of the Simple Mediated Call. the format field indicates that the location is an Entry Descriptor 47227. KOS Call microcode continues by saving the location of Entry Descriptor 47227 and creating a new 4 -666- Mediated Frame 46947 on current MAS Object 46703 and a new Ordinary SS Frame 10510 on SS Object 10336 for called Procedure 602- As KOS Call microcode does so, it sets Fields 46917 and 46919 in Mediated Frame Header 10414 and Fields 47109 and 47111 in Ordinary SS Frame Header 10514 to the values required by the addition of frames to MAS Object 46703 and SS Object 10336.

New Mediated Frame 46947 is now ready for Linkage Pointers 10416 to the actual arguments used in the Call, 1 0 so KOS Call microcode returns to S-interpreter Call microcode, which parses the SIN to obtain the literal specifying the number of arguments and saves the literal value. S-interpreter Call microcode then parses each argument Name, evaluates it, and places the resulting value in Linkage Pointers Section 10416. When Linkage Pointers Section 10416 is complete. S-Interpreter Call Microcode calculates the new location of FP from the location of the top of Linkage Pointers Section 10416 and |l tplaces a pointer for the location in the FU 10120 register reserved for FP. At this time. S-interpreter I II Call microcode also places the new location of the top of the stack in Stack Top Offset Field 46807.

*S-interpreter Call microcode then invokes KOS Call microcode to place the value of the literal specifying fWH 25 the number of arguments in MAS Frame Field 46929. to I calculate the new value of FHP 46702 and place it in the FU 10120 register reserved for that value, and finally to obtain the state necessary to execute called Procedure I i I -667- 602 from called Procedure 602's Entry Descriptor 47227 and PED 30303. As previously stated, S-interpreter Call microcode saved the location of Entry Descriptor 47227.

Using this location, KOS Call Microcode obtains the size of the storage required for local data from Field 47233 and adds that amount of storage to the new MAS Frame 46709- Then KOS Call Microcode uses Field 47231 to locate PED 30303 for Procedure 602. PED 30303 contains the remainder of the necessary information about F"t" 1 0 Procedure 602. and KOS Call microcode copies the location of PED 30303 into PED Pointer Field 46933 and then copies the values of K Field. 30305. Last Name Field 30307, NTP Field 30311. and PBP Field 30315 into the relevant registers in FU 10120- KOS Call microcode next translates the pointer in SIP Field 30309 into a dialect number as explained in Chapter 3, and places it in Sregister RDIAL 24212 of FU 10220 and thereupon derives SDP by resolving the pointer in SDPP Field 30313 and a pointer to SEB 46864 by resolving the pointer in SEPP 20 Field 30316. Having performed these operations, KOS Call microcode returns to S-interpreter Call microcode, which finishes the Call by obtaining a new PC, that is, resetting registers in I-stream Reader 27001 in FU 10120 ;so that the next SIN to be fetched will be the first SIN of called Procedure 602. S-interpreter Call microcode obtains the information required to change PC from Field 47229 in Entry Descriptor 47227 which contains the offset of the first SIN of called Procedure 602 from PBP.

-668- In the present embodiment, some FU 10120 state produced by the Mediated Call SIN is retained on SS 504 throughout the duration of Procedure 602's invocation.

The saved state allows Process 610 to reattempt the Mediated Call if the Call fails before the called Procedure 602 begins executing. When a Mediated Return SIN is executed, it resumes execution on the retained state from the CALL SINT. The Mediated Return is much simpler than the Call. Since all of the information i 10 required to resume execution of the invocation invocation which perfomed the Call is contained in Macrostate 10516 in the calling invocation's SS Frame 47003. Return need only pop the called invocation's frames from current MAS Object 46703 and SS Object 10336. copy Macrostate 10516 47123 from the calling invocation's SS Frame 47003 into the proper FU 10120 registers, translate SIP Value 47141 into a dialect number, and resume executing the calling invocation. The pop operation involves nothing more than Supdating those pointers in MAS Object 46703 and SS Object 20 10336 which pointed to locations in the old topmost frame so that they now point to equivalent locations in the new topmost frame.

c. InLoca.tina .of.PrQc.edures-60-2 i R e a u i'r i n -s g E f s -4 6 8 6 4 t 2 -7 l 25 468-;-469- 470 If a Procedure 602 requires a SEB 46864. this fact is indicated by Flag Field 47237 in Procedure 602's Entry Descriptor 47227. FED 30303 for such a Procedure 602 contains SEPP Field 47225. whose value is a non- -669resolvable pointer- The manner in which a SEB 46864 is created for Procedure 602 and SEPP field 47225 is translated into SEP, a pointer which contains the location of SEB 46864 and is saved as part of the invocation's macrostate on SS 10336. is similar to the J manner in which a Static Data Block 46863 is created and the non-resolvable pointer contained in SDPP field 47225 is translated into SDP. The first time that a Procedure ret* 602 requiring a SEB 46864 is invoked on a MAS Object 46703. a SEB 46864 is created for the Procedure 602 and an AATE 46857 is created which associates the nonresolvable pointer in SEPP field 47225 and the location of SEB 46864. That location is the value of SEP when the procedure is executing on MAS object 46703. On subsequent invocations of Procedure 602. AATE 46857 serves to translate the value in SEPP field 47225 into

SEP.

d.d. Crxss-P-cedu-re-bijec t-fc- z i (.Fig- 270. 46-8-" 469-;-470-- -47-1, 4221 A Mediated Call which invokes an External Procedure L* 602 is called a Cross-Procedure Object Call. As previously mentioned. KOS Call microcode assumes that any time the Name representing the called Procedure 602 in a Mediated Call SIN resolves to the location of a Gate that the Call is to an External Procedure 602- As long as newly-called External Procedure 602 has the same DOE as calling Procedure 602. Cross-Procedure Object Calls differ from the Simple Mediated Call only in the manner -670in which called Procedure 602's Entry Descriptor 47227 is located. Once KOS Call microcode has determined as descr-ibed above that a Mediated Call is a Cross-Procedure Object Call. it must next determine whether it is a Cross-Domain Call. To do so, KOS Call microcode compares the DOE Attribute of called Procedure 602's Procedure Object 608 with the domain component of the current subject. KOS Call microcode uses Procedure Object 608's AON 41304 to obtain Procedure Object 608's DOE from Field 10 41521 of its AOTE 41306. and it uses the ASN for the current subject, stored in an FU 10120 register, to l t °obtain the current subject's domain component from AST 10914. If the DOE and the current subject's domain component differ. the Call is a Cross-domain Call, described below; otherwise, the Call locates the Gate 47205 or 47206 specified by the evaluated Name for called Procedure 602 in its Procedure Object 608. If the Gate is a Local Gate 47205. the Call uses Entry Descriptor ,Offset Field 47207 to locate Entry Descriptor 47227 t 20 belonging to Called Procedure 602 and then proceeds as described in the discussion of a Simple Mediated Call.

r. If the Gate is a Link Gate 47206, KOS Call microcode "q obtains the pointer corresponding to Link Gate. 47206 from Binder Area 47245 and resolves it to obtain a pointer to another Gate 47205 or 47206. which KOS Call microcode t uses to repeat the External Procedure 602 call described above. The repetitions continue until the newly-located gate is a Local Gate 47205. whereupon Call proceeds as described for Simple Mediated Calls.

-671e.e. Cross-domain-Ca1l-i 2 48_. 418- 468. 469-. -470-.-4-7- If a called Procedure 602's Procedure Object 608 has a DOE attribute differing from that of calling Procedure 602's Procedure Object 608. the Call is a Cross-domain Call. The means by which KOS Call microcode determines that a Mediated Call is a Cross-Domain Call nave previously been described; If the Call is a Cross-Domain Call. KOS Call microcode must inactivate MAS Object 46703 for the domain from which the Call is made. perform trojan horse argument checks, switch subjects, place a Cross-domain Frame 47039 on SS object 10336, and locate and activate MAS Object 46703 for the new domain before it can make a Mediated Frame 46947 on new MAS Object 46703 and continue as described in the discussion of a Simple Mediated Call.

Cross-domain Call microcode first inactivates the t, current MAS Object 46703 by setting Domain Active Flag 20 46804 to FALSE. The next step is the trojan horse argument checks. In order to perform trojan horse argument checks, Cross-domain Call must have pointers to the actual arguments used in the cross-domain invocation. Consequently. Cross-domain Call first continues like a non-cross-domain Call: it creates a Mediated Frame Header 10414 on old MAS Object 46703 and returns to S-interpreter microcode. which evaluates the Names of the actual arguments, and places the pointers in 0 44 i 1 -,u -672- Linkage Pointers 10416 above Mediated Frame Header 10414. However. the macrostate for the invocation performing the call was placed on SS Object 10336 before Mediated Frame Header 10414 and Linkage Pointers 10416 were placed on old MAS Object 46703. Consequently. when calling Procedure 602 resumes execution after a Return, it will resume on MAS Frame 46709 preceding the one built by Cross-domain Call microcode.

Once the pointers to the actual arguments are available, Cross-domain Call Microcode performs the trojan horse check. As described in the discussion of Procedure Object 608 and illustrated in Figure 472. the information required to perform the check is contained in AIA 10352. Each Local Gate 47205 in Procedure Object 608 has an AIAE 47245. each formal argument in Local Gate 47205's prccedure has an entry in AIAE 47245's AMA 47251.

and the formal argument's AMAE 47253 indicates what kind of access to the formal argument's actual argument is required in called Procedure 602.

Field AIA OFF 47201 contains the location of AIA 10352 in Procedure Object 608. and using this information and Local Gate 47205's offset in Procedure Object 608, Cross-domain Call microcode locates AIAE 47245 for Local Gate 47205. The first two fields in AIAE 47245 contain the minimum number of arguments in the invocation and the maximum number of arguments. Cross-domain Call microcode checks whether the number of actual arguments falls between these values. If it does, Cross-domain Call microcode begins checking the access allowed individual -673arguments. For each argument pointer. Cross-domain Call microcode calls LAR microcode to obtain the current AON 41304 for the pointer's UID and uses AON 41304 and the ASN for Process 610's current subject the caller's subject) to locate an entry in either APAM 10918 or ANPAT 10920. depending on whether the argument's AIAE specifies primitive access (47255) or extended access (47257) respectively. If tihe information from APAM 10918 or ANPAT 10920 confirms that Process 610's current subject S 10 has the right to access the argument in the manner required in called Procedure 602. the Trojan Horse microcode goes on to the next argument. If the current subject has the required access to all arguments, the trojan horse check succeeds and the Cross-domain Call continues. Otherwise, it fails and Cross-domain Call performs a microcode-to-software Call as explained below.

Next, Cross-domain Call microcode places Cross domain State 10513 on sS Object 10336. As explained in the discussion of SS object 10336. Cross-domain State 10513 20 contains the information required to return to the caller's frame on former MAS Object 46703. Having done this, Cross-domain Call microcode changes subjects.

Using the current subject's ASN. Cross-Domain Call SA 1microcode obtains the current subject from AST 10914.

replaces the subject's domain component with DOE Attribute 41225 for called Procedure 602's Procedure Object 608. and uses AST 10914 to translate the new subject thus obtained into a new ASN. That ASN then is placed in the appropriate FU 10120 register.

-674- After the subject has been changed, Cross-domain Call microcode uses Domain Table 41801 to translate the DOE of called Procedure 602 into a domain number. Cross-domain Call microcode then uses the domain number as an index into Array of MAS AONs 46211 in VPSB 614 for Virtual Processor 612 belonging to Process 610 making the crossdomain call. The entry corresponding to the domain number contains AON 41304 of MAS Object 46703 for that n domain.

Having located the proper MAS Object 46703, Crossdomain Call microcode uses STO field 46807 in MAS Header 10410 belonging to the new domains MAS Object 46703 to locate the top of the last MAS Frame 46709. It then saves the value of FHP 46702 used in the preceding invocation in a FU 10120 register. adds a Mediated Frame l Header 10414 to the top of MAS Object 46703, and c calculates a new FHP 46702 which points to new Mediated Frame Header 10414. KOS Cross-Domain Call microcode then places the old value of FHP 46702 in FHP Value Field 47151 of SS Object 10336 and the old value of STO 46704 (pointing to the top of the last complete MAS Frame 46709 on previous MAS Object 46703) in Field 47153 of Cross- S Domain State 10513 and fills in.Mediated Frame Header 10414 fields as follows: Result of Cross-domain Call Field 46903 is set to TRUE, Previous Frame Offset Field 46917 is set to 0, and Dynamic Back Pointer Field 46931 is set to the saved value of FHP 46702. Dynamic Back Pointer Field 46931 thus points to the header of the -675topmost Mediated Frame 46947 on the previous MAS Object 46703. The values of the remaining fields are copied from Mediated Frame Header 10414 which Cross-Domain Call created on previous MAS Object 46703.

Cross-domain Call microcode next copies the argument pointers for the formal arguments from the top of previous MAS Object 46703 to new Mediated Frame 46947 and calculates FP. Cross-domain Call Microcode finishes by returning to S-interpreter Call microcode, Which t0 completes the Call as described for Simple Mediated Calls.

Except for the work involved in transferring to a new MAS Object 46703. Cross-domain Return is like other Returns from Mediated Calls. Old FHP 46701. from Field 47151 of Cross-Domain State 10513. and old STO 46704.

from Field 47153 of Cross-domain State, are placed in FU 10120 registers. Then the frames belonging to the invocation that is ending are popped off of SS Object 10336 and off of MAS Object 46703 belonging to the domain of called Procedure 602. and MAS Object 46703 is inactivated by setting Domain Active Flag 46804 to FALSE. Then KOS Cross-domain Return microcode uses old FHP 46701 and old STO 46704 to locate MAS Object 46703 being returned to and the topmost Mediated Frame 46947 on that MAS Object 46703. MAS Object 46703 being returned

J--

to is activated, and finally, the contents of Macrostate 10516 belonging to the invocation being returned to are placed in the appropriate registers of FU 10120 and execution of the invocation resumes.

1 k a -676f f. Failed Crs s-DomainlCa--s-fs-i 2720;'.-4 68--46-47-A 4L 472) A Cross-Domain Call as described above may fail at several points between the time that the calling 3 invocation begins the call and called Procedure 602 begins executing. On failure. Cross-Domain Call microcode performs a microcode-to-software Call. KOS Procedures 602 invoked by this Call may remedy the reason for the Cross Domain Call's failure and reattempt the Cross-domain Call. This is possible because the implementation of Oross Domain Call in CS 10110 saves sufficient FU 10120 state to allow Process 610 executing the Cross-Domain Call to return to the invocation and the Mediated Call SIN from which the Cross-Domain Call began. On failure, the invocation's MAS Frame 46709 may be located from the values of STO Field 47153 and FHP Field 47151 in Cross-Domain State 10513. and the Mediated Call SIN may be located by using information saved in FU 10120 state.

6. ig bohood--Ga-1s-----4 6-46-9- 47 21 As previously mentioned. Procedures 602 called via Neighborhood Calls must have the same PED 30303 as calling Procedure 602. The only macrostate values which are not part of PED 30303 are PC and FP; consequently.

Neighborhood Call need only save PC and FP of the invocation performing the call and calculate these values for the new invocation. In addition, Neighborhood Call saves STO 46704 in order to make it easier to locate the top of the previous invocation's Neighborhood Frame 46947. Neighborhood Return simply restores the saved sr -677values. Since the macrostate values copied from or obtained via PED 30303 do not change during the sequence of invocations, and therefore need not be saved on SS Object 10336. Neighborhood Calls do not have SS Frames 47003.

The invention may be embodied in yet other specific forms without departing from the spirit or essential characteristics thereof. Thus, the present embodiments are to be considered in all respects as illustrative and 10 not restrictive, the scope of the invention being indicated by the appended claims rather than by the foregoing description, and all changes which come within this meaning and range of equivalency of the claims are therefore intended to be embraced therein.

I

it (t It I t t ti

I.

I It I I I &Itf It I It 'i 1, 7/ V L

Claims (13)

1. 2. A digital computer system of claim 1 wherein said memory operation specifier further includes a length ;A specifier specifying the length of said item at the location Sspecified by said logical address; and MAP 678 WV' MJP 678 -t. said memory operation means performs said memory opera- tion specified by said memory command on said item at said location specified by said logical address having the length specified by said length specifier.
2. 3. A digital computer system of claim 2 wherein said length specifier specifies a length in bits.
3. 4. A digital computer system of claim 1 wherein said offset is a bit-granular offset and may specify any bit of the items associated with the object specified in said logical address. A digital computer system of claim 1 wherein said *4* memory organization means identifies each said object with a single unique identifier and said single unique identifier never identifies any other said object. rrt
4. 6. A digital computer system of claim 1 wherein said n. unique identifier contains 80 bits.
5. 7. A digital computer system of claim 1 wherein said memory organization means includes means for organizing said memory means into a plurality of logical allocation units, I!t* each said object being associated with one of said logical allocation units, and said unique identifier further includes r ,a logical allocation unit identifier specifying the S logical allocation unit associated with the object identified by said unique identifier, and A an object serial number identifying said object in said associated logical allocation unit.
6. 8. A digital computer system of claim 7 wherein said memory organization means identifies each said logical allocation unit by a single logical allocation unit identifier and said single logical allocation unit identifier never identifies any other said logical allocation unit, and 39 said memory organization means identifies each object BL -679 ~>4 -7b S- associated with said logical allocation unit by a single object serial number and said object serial number never identifies any other said object associated with said logical allocation unit.
7. 9. A digital computer system of claim 7 wherein said logical allocation unit identifier further includes a logical allocation unit group number specifying a logical allocation group including a plurality of said.logical allocation units, and a logical allocation unit serial number identifying one of the logical allocation units of said plurality of logical allocation units. A digital computer system of claim 9 wherein said memory organization means identifies each of said S pluralityof logical allocation unit groups by a single logical allocation unit group number and said single logical allocation unit group number never identifies any other said logical allocation unit group; and 2C said memory organization means identifies each logical allocation unit associated with the logical allocation unit group identified by said logical allocation unit group number by a single logical allocation unit serial number and said single logical allocation unit serial number never identifies S any other logical allocation unit in the logical allocation unit group specified by said logical allocation unit group number. t
8. 11. A digital computer system of claim 9 wherein said logical allocation unit group number contains 9 <.bits; said logical allocation unit serial number contains 8 bits; i said object serial number contains 48 bits; and said offset identifies a single bit which may be any bit in the items associated with the object identified by said unique identifier. 39 12. A digital computer system of claim 1 wherein said BL 680 J r,3 universal identifiers do not specify locations of objects relative to other said objects.
9. 13. A digital computer system of claim 1 wherein certain ones of said subjects, including said current subject, are active subjects for which said processor means is or shortly will be executing said instructions; said memory operations means includes means for temporarily associating each said active subject with a different active subject number, said current subject being represented in said memory operation means by the active subject number currently associated therewith and said memory operations means using the active subject number currently associated with said current subject to determine whether said current subject is one of the set of subjects identified for the object specified in said memory operation specifier.
10. 14. A digital computer system of claim 13 wherein said memory operation means further i-ncludes means accessible to said processor means for storing the active subject number currently associated with said current subject. A digital computer system of claim 1 wherein said memory organization means further identifies for each of said selected set of subjects a selected set of said memory operations which said selected set of subjects is authorized to perform on the items locatable by a said object; and said memory operation means further performs the memory operation specified in said memory operation specifier only if said specified memory operation is identified as one of the selected set of memory operations which said current subject is authorized to perform on the items locatable by the object specified in said memory operation specifier.
11. 16. A digital computer system of claim 15 wherein said 39 memory operation means further includes BI; 681 r protection cache means for encaching information specifying said memory operations which a current subject is authorized to perform on the objects specified in the memory operation specifiers produced in response to said instructions, said protection cache means receiving said memory operation specifier and responding thereto by producing an access violation signal when said encached information indicates that said current subject may not perform the memory operation specified by said memory operation specifier on the object specified by said memory operation specifier; and means responsive to said access violation signal for inhibiting the performance of said specified memory operation.
12. 17. A digital computer system of claim 1 wherein ,Etc said instructions are contained in procedures of said items and said processor means responds to said instructions S only while executing one of said procedures; the operations performed by said processor means further S include 2 a call operation for suspending a current execution of one of said procedures and commencing another execution as the S. current execution, and a return operation for terminating said current execution S t and resuming the execution which was suspended to commence the terminated current execution; and said current subject changes as a consequence of said call operation and said return operation. U t V C~A SDATED Augut, 9 B'-IL PSCRMNDE 6iT8PATfICK r I- 39 BL 682 A.j i| r 9999 9 *999 9 9 99 *9 4 4 999 9 *999., 4 9 '4 .9 4 999s' 9*9* 90 9* S *9 9 9 9 .9 9 9 4 4
13. 18. A digital computer system according to claim 1 substantially as herein described with reference to the accompanying drawings. DATED: 27 November, 1989. PHILLIPS ORMONDE FITZPATRICK Attorneys for: DATA GENERAL CORPORATION M4JP 683